The RISC-V Instruction Set Manual: Volume II: Privileged Architecture

Andrew waterman, waterman@eecs.berkeley.edu; Krste Asanović, krste@berkeley.edu; John Hauser, jh.riscv@jhauser.us, SiFive Inc., CS Division, EECS Department, University of California, Berkeley
version 20211203, Revised 20230731

Table of Contents

• The RISC-V Instruction Set Manual: Volume II: Privileged Architecture
• Preface
• 1. Introduction
  • 1.1. RISC-V Privileged Software Stack Terminology
  • 1.2. Privilege Levels
  • 1.3. Debug Mode
• 2. Control and Status Registers (CSRs)
  • 2.1. CSR Address Mapping Conventions
  • 2.2. CSR Listing
  • 2.3. CSR Field Specifications
  • 2.3.1. Reserved Writes Preserve Values, Reads Ignore Values (WPRI)
  • 2.3.2. Write/Read Only Legal Values (WLRL)
  • 2.3.3. Write Any Values, Reads Legal Values (WARL)
  • 2.4. CSR Field Modulation
  • 2.5. Implicit Reads of CSRs
  • 2.6. CSR Width Modulation
  • 2.7. Explicit Accesses to CSRs Wider than XLEN
• 3. Machine-Level ISA, Version 1.12
  • 3.1. Machine-Level CSRs
  • 3.1.1. Machine ISA Register misa
  • 3.1.2. Machine Vendor ID Register mvendorid
  • 3.1.3. Machine Architecture ID Register marchid
  • 3.1.4. Machine Implementation ID Register mimpid
  • 3.1.5. Hart ID Register mhartid
  • 3.1.6. Machine Status Registers (mstatus and mstatush)
    • 3.1.6.1. Privilege and Global Interrupt-Enable Stack in mstatus register
    • 3.1.6.2. Base ISA Control in mstatus Register
    • 3.1.6.3. Memory Privilege in mstatus Register
    • 3.1.6.4. Endianness Control in mstatus and mstatush Registers
    • 3.1.6.5. Virtualization Support in mstatus Register
    • 3.1.6.6. Extension Context Status in mstatus Register
  • 3.1.7. Machine Trap-Vector Base-Address Register (mtvec)
  • 3.1.8. Machine Trap Delegation Registers (medeleg and mideleg)
  • 3.1.9. Machine Interrupt Registers (mip and mie)
  • 3.1.10. Hardware Performance Monitor
  • 3.1.11. Machine Counter-Enable Register (mcounteren)
  • 3.1.12. Machine Counter-Inhibit CSR (mcountinhibit)
  • 3.1.13. Machine Scratch Register (mscratch)
  • 3.1.14. Machine Exception Program Counter (mepc)
  • 3.1.15. Machine Cause Register (mcause)
  • 3.1.16. Machine Trap Value Register (mtval)
  • 3.1.17. Machine Configuration Pointer Register (mconfigptr)
  • 3.1.18. Machine Environment Configuration Register (menvcfg)
  • 3.1.19. Machine Security Configuration Register (mseccfg)
  • 3.2. Machine-Level Memory-Mapped Registers
  • 3.2.1. Machine Timer Registers (mtime and mtimecmp)
  • 3.3. Machine-Mode Privileged Instructions
  • 3.3.1. Environment Call and Breakpoint
  • 3.3.2. Trap-Return Instructions
  • 3.3.3. Wait for Interrupt
  • 3.3.4. Custom SYSTEM Instructions
  • 3.4. Reset
  • 3.5. Non-Maskable Interrupts
  • 3.6. Physical Memory Attributes
  • 3.6.1. Main Memory versus I/O versus Vacant Regions
  • 3.6.2. Supported Access Type PMAs
  • 3.6.3. Atomicity PMAs
    • 3.6.3.1. AMO PMA
    • 3.6.3.2. Reservability PMA
    • 3.6.3.3. Alignment
  • 3.6.4. Memory-Ordering PMAs
  • 3.6.5. Coherence and Cacheability PMAs
  • 3.6.6. Idempotency PMAs
  • 3.7. Physical Memory Protection
  • 3.7.1. Physical Memory Protection CSRs
    • 3.7.1.1. Address Matching
    • 3.7.1.2. Locking and Privilege Mode
    • 3.7.1.3. Priority and Matching Logic
  • 3.7.2. Physical Memory Protection and Paging
• 4. "Smrnmi" Standard Extension for Resumable Non-Maskable Interrupts, Version 0.4
  • 4.1. RNMI Interrupt Signals
  • 4.2. RNMI Handler Addresses
  • 4.3. RNMI CSRs
  • 4.4. MNRET Instruction
  • 4.5. RNMI Operation
• 5. Supervisor-Level ISA, Version 1.12
  • 5.1. Supervisor CSRs
  • 5.1.1. Supervisor Status Register (sstatus)
    • 5.1.1.1. Base ISA Control in sstatus Register
    • 5.1.1.2. Memory Privilege in sstatus Register
    • 5.1.1.3. Endianness Control in sstatus Register
  • 5.1.2. Supervisor Trap Vector Base Address Register (stvec)
  • 5.1.3. Supervisor Interrupt Registers (sip and sie)
  • 5.1.4. Supervisor Timers and Performance Counters
  • 5.1.5. Counter-Enable Register (scounteren)
  • 5.1.6. Supervisor Scratch Register (sscratch)
  • 5.1.7. Supervisor Exception Program Counter (sepc)
  • 5.1.8. Supervisor Cause Register (scause)
  • 5.1.9. Supervisor Trap Value (stval) Register
  • 5.1.10. Supervisor Environment Configuration Register (senvcfg)
  • 5.1.11. Supervisor Address Translation and Protection (satp) Register
  • 5.2. Supervisor Instructions
  • 5.2.1. Supervisor Memory-Management Fence Instruction
  • 5.3. Sv32: Page-Based 32-bit Virtual-Memory Systems
  • 5.3.1. Addressing and Memory Protection
  • 5.3.2. Virtual Address Translation Process
  • 5.4. Sv39: Page-Based 39-bit Virtual-Memory System
  • 5.4.1. Addressing and Memory Protection
  • 5.5. Sv48: Page-Based 48-bit Virtual-Memory System
  • 5.5.1. Addressing and Memory Protection
  • 5.6. Sv57: Page-Based 57-bit Virtual-Memory System
  • 5.6.1. Addressing and Memory Protection
• 6. "Svnapot" Standard Extension for NAPOT Translation Contiguity, Version 1.0
• 7. "Svpbmt" Standard Extension for Page-Based Memory Types, Version 1.0
• 8. "Svinval" Standard Extension for Fine-Grained Address-Translation Cache Invalidation, Version 1.0
• 9. Hypervisor Extension, Version 1.0
  • 9.1. Privilege Modes
  • 9.2. Hypervisor and Virtual Supervisor CSRs
  • 9.2.1. Hypervisor Status Register (hstatus)
  • 9.2.2. Hypervisor Trap Delegation Registers (hedeleg and hideleg)
  • 9.2.3. Hypervisor Interrupt Registers (hvip, hip, and hie)
  • 9.2.4. Hypervisor Guest External Interrupt Registers (hgeip and hgeie)
  • 9.2.5. Hypervisor Environment Configuration Register (henvcfg)
  • 9.2.6. Hypervisor Counter-Enable Register (hcounteren)
  • 9.2.7. Hypervisor Time Delta Register (htimedelta)
  • 9.2.8. Hypervisor Trap Value Register (htval)
  • 9.2.9. Hypervisor Trap Instruction Register (htinst)
  • 9.2.10. Hypervisor Guest Address Translation and Protection Register (hgatp)
  • 9.2.11. Virtual Supervisor Status Register (vsstatus)
  • 9.2.12. Virtual Supervisor Interrupt Registers (vsip and vsie)
  • 9.2.13. Virtual Supervisor Trap Vector Base Address Register (vstvec)
  • 9.2.14. Virtual Supervisor Scratch Register (vsscratch)
  • 9.2.15. Virtual Supervisor Exception Program Counter (vsepc)
  • 9.2.16. Virtual Supervisor Cause Register (vscause)
  • 9.2.17. Virtual Supervisor Trap Value Register (vstval)
  • 9.2.18. Virtual Supervisor Address Translation and Protection Register (vsatp)
  • 9.3. Hypervisor Instructions
  • 9.3.1. Hypervisor Virtual-Machine Load and Store Instructions
  • 9.3.2. Hypervisor Memory-Management Fence Instructions
  • 9.4. Machine-Level CSRs
  • 9.4.1. Machine Status Registers (mstatus and mstatush)
  • 9.4.2. Machine Interrupt Delegation Register (mideleg)
  • 9.4.3. Machine Interrupt Registers (mip and mie)
  • 9.4.4. Machine Second Trap Value Register (mtval2)
  • 9.4.5. Machine Trap Instruction Register (mtinst)
  • 9.5. Two-Stage Address Translation
  • 9.5.1. Guest Physical Address Translation
  • 9.5.2. Guest-Page Faults
  • 9.5.3. Memory-Management Fences
  • 9.6. Traps
  • 9.6.1. Trap Cause Codes
  • 9.6.2. Trap Entry
  • 9.6.3. Transformed Instruction or Pseudoinstruction for mtinst or htinst
  • 9.6.4. Trap Return
• 10. RISC-V Privileged Instruction Set Listings
• 11. History
  • 11.1. Research Funding at UC Berkeley
• Bibliography

Contributors to all versions of the spec in alphabetical order (please contact editors to suggest corrections): Krste Asanović, Peter Ashenden, Rimas Avižienis, Jacob Bachmeyer, Allen J. Baum, Jonathan Behrens, Paolo Bonzini, Ruslan Bukin, Christopher Celio, Chuanhua Chang, David Chisnall, Anthony Coulter, Palmer Dabbelt, Monte Dalrymple, Paul Donahue, Greg Favor, Dennis Ferguson, Marc Gauthier, Andy Glew, Gary Guo, Mike Frysinger, John Hauser, David Horner, Olof Johansson, David Kruckemyer, Yunsup Lee, Daniel Lustig, Andrew Lutomirski, Prashanth Mundkur, Jonathan Neuschäfer, Rishiyur Nikhil, Stefan O’Rear, Albert Ou, John Ousterhout, David Patterson, Dmitri Pavlov, Kade Phillips, Josh Scheid, Colin Schmidt, Michael Taylor, Wesley Terpstra, Matt Thomas, Tommy Thorn, Ray VanDeWalker, Megan Wachs, Steve Wallach, Andrew Waterman, Claire Wolf, and Reinoud Zandijk..

This document is released under a Creative Commons Attribution 4.0 International License.

This document is a derivative of the RISC-V privileged specification version 1.9.1 released under following license: ©2010-2017 Andrew Waterman, Yunsup Lee, Rimas Avižienis, David Patterson, Krste Asanović. Creative Commons Attribution 4.0 International License.

Please cite as: “The RISC-V Instruction Set Manual, Volume II: Privileged Architecture, Document Version 20211203”, Editors Andrew Waterman, Krste Asanović, and John Hauser, RISC-V International, December 2021.

Preface

This document describes the RISC-V privileged architecture. This release, version 20211203, contains the following versions of the RISC-V ISA modules:

Module	Version	Status
Machine ISA Smrnmi Extension Supervisor ISA Svnapot Extension Svpbmt Extension Svinval Extension Hypervisor ISA	1.13 0.1 1.12 1.0 1.0 1.0 1.0	Draft Draft Ratified Ratified Ratified Ratified Ratified

The following compatible changes have been made to the Machine ISA since version 1.12:

• Defined the misa.V field to reflect that the V extension has been implemented.

• Clarified semantics of explicit accesses to CSRs wider than XLEN bits.

• Clarified that MXLEN≥SXLEN, and added the constraint that SXLEN≥UXLEN.

• Clarified that WFI is not a HINT instruction.

• Clarified ordering rules when PBMT=IO is used on main-memory regions.

Preface to Version 20211203

This document describes the RISC-V privileged architecture. This release, version 20211203, contains the following versions of the RISC-V ISA modules:

Module	Version	Status
Machine ISA Supervisor ISA Svnapot Extension Svpbmt Extension Svinval Extension Hypervisor ISA	1.12 1.12 1.0 1.0 1.0 1.0	Ratified Ratified Ratified Ratified Ratified Ratified

The following changes have been made since version 1.11, which, while not strictly backwards compatible, are not anticipated to cause software portability problems in practice:

• Changed MRET and SRET to clear mstatus.MPRV when leaving M-mode.

• Reserved additional satp patterns for future use.

• Stated that the scause Exception Code field must implement bits 4–0 at minimum.

• Relaxed I/O regions have been specified to follow RVWMO. The previous specification implied that PPO rules other than fences and acquire/release annotations did not apply.

• Constrained the LR/SC reservation set size and shape when using page-based virtual memory.

• PMP changes require an SFENCE.VMA on any hart that implements page-based virtual memory, even if VM is not currently enabled.

• Allowed for speculative updates of page table entry A bits.

• Clarify that if the address-translation algorithm non-speculatively reaches a PTE in which a bit reserved for future standard use is set, a page-fault exception must be raised.

Additionally, the following compatible changes have been made since version 1.11:

• Removed the N extension.

• Defined the mandatory RV32-only CSR mstatush, which contains most of the same fields as the upper 32 bits of RV64’s mstatus.

• Defined the mandatory CSR mconfigptr, which if nonzero contains the address of a configuration data structure.

• Defined optional mseccfg and mseccfgh CSRs, which control the machine’s security configuration.

• Defined menvcfg, henvcfg, and senvcfg CSRs (and RV32-only menvcfgh and henvcfgh CSRs), which control various characteristics of the execution environment.

• Designated part of SYSTEM major opcode for custom use.

• Permitted the unconditional delegation of less-privileged interrupts.

• Added optional big-endian and bi-endian support.

• Made priority of load/store/AMO address-misaligned exceptions implementation-defined relative to load/store/AMO page-fault and access-fault exceptions.

• PMP reset values are now platform-defined.

• An additional 48 optional PMP registers have been defined.

• Slightly relaxed the atomicity requirement for A and D bit updates performed by the implementation.

• Clarify the architectural behavior of address-translation caches

• Added Sv57 and Sv57x4 address translation modes.

• Software breakpoint exceptions are permitted to write either 0 or the pc to xtval.

• Clarified that bare S-mode need not support the SFENCE.VMA instruction.

• Specified relaxed constraints for implicit reads of non-idempotent regions.

• Added the Svnapot Standard Extension, along with the N bit in Sv39, Sv48, and Sv57 PTEs.

• Added the Svpbmt Standard Extension, along with the PBMT bits in Sv39, Sv48, and Sv57 PTEs.

• Added the Svinval Standard Extension and associated instructions.

Finally, the hypervisor architecture proposal has been extensively revised.

Preface to Version 1.11

This is version 1.11 of the RISC-V privileged architecture. The document contains the following versions of the RISC-V ISA modules:

Module	Version	Status
Machine ISA Supervisor ISA Hypervisor ISA	1.11 1.11 0.3	Ratified Ratified Draft

Changes from version 1.10 include:

• Moved Machine and Supervisor spec to Ratified status.

• Improvements to the description and commentary.

• Added a draft proposal for a hypervisor extension.

• Specified which interrupt sources are reserved for standard use.

• Allocated some synchronous exception causes for custom use.

• Specified the priority ordering of synchronous exceptions.

• Added specification that xRET instructions may, but are not required to, clear LR reservations if A extension present.

• The virtual-memory system no longer permits supervisor mode to execute instructions from user pages, regardless of the SUM setting.

• Clarified that ASIDs are private to a hart, and added commentary about the possibility of a future global-ASID extension.

• SFENCE.VMA semantics have been clarified.

• Made the mstatus.MPP field WARL, rather than WLRL.

• Made the unused xip fields WPRI, rather than WIRI.

• Made the unused misa fields WARL, rather than WIRI.

• Made the unused pmpaddr and pmpcfg fields WARL, rather than WIRI.

• Required all harts in a system to employ the same PTE-update scheme as each other.

• Rectified an editing error that misdescribed the mechanism by which mstatus.xIE is written upon an exception.

• Described scheme for emulating misaligned AMOs.

• Specified the behavior of the misa and xepc registers in systems with variable IALIGN.

• Specified the behavior of writing self-contradictory values to the misa register.

• Defined the mcountinhibit CSR, which stops performance counters from incrementing to reduce energy consumption.

• Specified semantics for PMP regions coarser than four bytes.

• Specified contents of CSRs across XLEN modification.

• Moved PLIC chapter into its own document.

Preface to Version 1.10

This is version 1.10 of the RISC-V privileged architecture proposal. Changes from version 1.9.1 include:

• The previous version of this document was released under a Creative Commons Attribution 4.0 International License by the original authors, and this and future versions of this document will be released under the same license.

• The explicit convention on shadow CSR addresses has been removed to reclaim CSR space. Shadow CSRs can still be added as needed.

• The mvendorid register now contains the JEDEC code of the core provider as opposed to a code supplied by the Foundation. This avoids redundancy and offloads work from the Foundation.

• The interrupt-enable stack discipline has been simplified.

• An optional mechanism to change the base ISA used by supervisor and user modes has been added to the mstatus CSR, and the field previously called Base in misa has been renamed to MXL for consistency.

• Clarified expected use of XS to summarize additional extension state status fields in mstatus.

• Optional vectored interrupt support has been added to the mtvec and stvec CSRs.

• The SEIP and UEIP bits in the mip CSR have been redefined to support software injection of external interrupts.

• The mbadaddr register has been subsumed by a more general mtval register that can now capture bad instruction bits on an illegal instruction fault to speed instruction emulation.

• The machine-mode base-and-bounds translation and protection schemes have been removed from the specification as part of moving the virtual memory configuration to sptbr (now satp). Some of the motivation for the base and bound schemes are now covered by the PMP registers, but space remains available in mstatus to add these back at a later date if deemed useful.

• In systems with only M-mode, or with both M-mode and U-mode but without U-mode trap support, the medeleg and mideleg registers now do not exist, whereas previously they returned zero.

• Virtual-memory page faults now have mcause values distinct from physical-memory access faults. Page-fault exceptions can now be delegated to S-mode without delegating exceptions generated by PMA and PMP checks.

• An optional physical-memory protection (PMP) scheme has been proposed.

• The supervisor virtual memory configuration has been moved from the mstatus register to the sptbr register. Accordingly, the sptbr register has been renamed to satp (Supervisor Address Translation and Protection) to reflect its broadened role.

• The SFENCE.VM instruction has been removed in favor of the improved SFENCE.VMA instruction.

• The mstatus bit MXR has been exposed to S-mode via sstatus.

• The polarity of the PUM bit in sstatus has been inverted to shorten code sequences involving MXR. The bit has been renamed to SUM.

• Hardware management of page-table entry Accessed and Dirty bits has been made optional; simpler implementations may trap to software to set them.

• The counter-enable scheme has changed, so that S-mode can control availability of counters to U-mode.

• H-mode has been removed, as we are focusing on recursive virtualization support in S-mode. The encoding space has been reserved and may be repurposed at a later date.

• A mechanism to improve virtualization performance by trapping S-mode virtual-memory management operations has been added.

• The Supervisor Binary Interface (SBI) chapter has been removed, so that it can be maintained as a separate specification.

Preface to Version 1.9.1

This is version 1.9.1 of the RISC-V privileged architecture proposal. Changes from version 1.9 include:

• Numerous additions and improvements to the commentary sections.

• Change configuration string proposal to be use a search process that supports various formats including Device Tree String and flattened Device Tree.

• Made misa optionally writable to support modifying base and supported ISA extensions. CSR address of misa changed.

• Added description of debug mode and debug CSRs.

• Added a hardware performance monitoring scheme. Simplified the handling of existing hardware counters, removing privileged versions of the counters and the corresponding delta registers.

• Fixed description of SPIE in presence of user-level interrupts.

1. Introduction

This document describes the RISC-V privileged architecture, which covers all aspects of RISC-V systems beyond the unprivileged ISA, including privileged instructions as well as additional functionality required for running operating systems and attaching external devices.

Commentary on our design decisions is formatted as in this paragraph, and can be skipped if the reader is only interested in the specification itself. We briefly note that the entire privileged-level design described in this document could be replaced with an entirely different privileged-level design without changing the unprivileged ISA, and possibly without even changing the ABI. In particular, this privileged specification was designed to run existing popular operating systems, and so embodies the conventional level-based protection model. Alternate privileged specifications could embody other more flexible protection-domain models. For simplicity of expression, the text is written as if this was the only possible privileged architecture.

1.1. RISC-V Privileged Software Stack Terminology

This section describes the terminology we use to describe components of the wide range of possible privileged software stacks for RISC-V.

Figure 1 shows some of the possible software stacks that can be supported by the RISC-V architecture. The left-hand side shows a simple system that supports only a single application running on an application execution environment (AEE). The application is coded to run with a particular application binary interface (ABI). The ABI includes the supported user-level ISA plus a set of ABI calls to interact with the AEE. The ABI hides details of the AEE from the application to allow greater flexibility in implementing the AEE. The same ABI could be implemented natively on multiple different host OSs, or could be supported by a user-mode emulation environment running on a machine with a different native ISA.

Our graphical convention represents abstract interfaces using black boxes with white text, to separate them from concrete instances of components implementing the interfaces.

Figure 1. Different implementation stacks supporting various forms of privileged execution.

The middle configuration shows a conventional operating system (OS) that can support multiprogrammed execution of multiple applications. Each application communicates over an ABI with the OS, which provides the AEE. Just as applications interface with an AEE via an ABI, RISC-V operating systems interface with a supervisor execution environment (SEE) via a supervisor binary interface (SBI). An SBI comprises the user-level and supervisor-level ISA together with a set of SBI function calls. Using a single SBI across all SEE implementations allows a single OS binary image to run on any SEE. The SEE can be a simple boot loader and BIOS-style IO system in a low-end hardware platform, or a hypervisor-provided virtual machine in a high-end server, or a thin translation layer over a host operating system in an architecture simulation environment.

Most supervisor-level ISA definitions do not separate the SBI from the execution environment and/or the hardware platform, complicating virtualization and bring-up of new hardware platforms.

The rightmost configuration shows a virtual machine monitor configuration where multiple multiprogrammed OSs are supported by a single hypervisor. Each OS communicates via an SBI with the hypervisor, which provides the SEE. The hypervisor communicates with the hypervisor execution environment (HEE) using a hypervisor binary interface (HBI), to isolate the hypervisor from details of the hardware platform.

The ABI, SBI, and HBI are still a work-in-progress, but we are now prioritizing support for Type-2 hypervisors where the SBI is provided recursively by an S-mode OS.

Hardware implementations of the RISC-V ISA will generally require additional features beyond the privileged ISA to support the various execution environments (AEE, SEE, or HEE).

1.2. Privilege Levels

At any time, a RISC-V hardware thread (hart) is running at some privilege level encoded as a mode in one or more CSRs (control and status registers). Three RISC-V privilege levels are currently defined as shown in Table 1.

Table 1. RISC-V privilege levels.

Level	Encoding	Name	Abbreviation
0 1 2 3	00 01 10 11	User/Application Supervisor Reserved Machine	U S   M

Table 1. RISC-V privilege levels. {#privlevels}

Privilege levels are used to provide protection between different components of the software stack, and attempts to perform operations not permitted by the current privilege mode will cause an exception to be raised. These exceptions will normally cause traps into an underlying execution environment.

In the description, we try to separate the privilege level for which code is written, from the privilege mode in which it runs, although the two are often tied. For example, a supervisor-level operating system can run in supervisor-mode on a system with three privilege modes, but can also run in user-mode under a classic virtual machine monitor on systems with two or more privilege modes. In both cases, the same supervisor-level operating system binary code can be used, coded to a supervisor-level SBI and hence expecting to be able to use supervisor-level privileged instructions and CSRs. When running a guest OS in user mode, all supervisor-level actions will be trapped and emulated by the SEE running in the higher-privilege level.

The machine level has the highest privileges and is the only mandatory privilege level for a RISC-V hardware platform. Code run in machine-mode (M-mode) is usually inherently trusted, as it has low-level access to the machine implementation. M-mode can be used to manage secure execution environments on RISC-V. User-mode (U-mode) and supervisor-mode (S-mode) are intended for conventional application and operating system usage respectively.

Each privilege level has a core set of privileged ISA extensions with optional extensions and variants. For example, machine-mode supports an optional standard extension for memory protection. Also, supervisor mode can be extended to support Type-2 hypervisor execution as described in Chapter 9.

Implementations might provide anywhere from 1 to 3 privilege modes trading off reduced isolation for lower implementation cost, as shown in Table 2.

Table 2. Supported combination of privilege modes.

Number of levels	Supported Modes	Intended Usage
1 2 3	M M, U M, S, U	Simple embedded systems Secure embedded systems Systems running Unix-like operating systems

Table 2. Supported combination of privilege modes. {#privcombs}

All hardware implementations must provide M-mode, as this is the only mode that has unfettered access to the whole machine. The simplest RISC-V implementations may provide only M-mode, though this will provide no protection against incorrect or malicious application code.

The lock feature of the optional PMP facility can provide some limited protection even with only M-mode implemented.

Many RISC-V implementations will also support at least user mode (U-mode) to protect the rest of the system from application code. Supervisor mode (S-mode) can be added to provide isolation between a supervisor-level operating system and the SEE.

A hart normally runs application code in U-mode until some trap (e.g., a supervisor call or a timer interrupt) forces a switch to a trap handler, which usually runs in a more privileged mode. The hart will then execute the trap handler, which will eventually resume execution at or after the original trapped instruction in U-mode. Traps that increase privilege level are termed vertical traps, while traps that remain at the same privilege level are termed horizontal traps. The RISC-V privileged architecture provides flexible routing of traps to different privilege layers.

Horizontal traps can be implemented as vertical traps that return control to a horizontal trap handler in the less-privileged mode.

1.3. Debug Mode

Implementations may also include a debug mode to support off-chip debugging and/or manufacturing test. Debug mode (D-mode) can be considered an additional privilege mode, with even more access than M-mode. The separate debug specification proposal describes operation of a RISC-V hart in debug mode. Debug mode reserves a few CSR addresses that are only accessible in D-mode, and may also reserve some portions of the physical address space on a platform.

2. Control and Status Registers (CSRs)

The SYSTEM major opcode is used to encode all privileged instructions in the RISC-V ISA. These can be divided into two main classes: those that atomically read-modify-write control and status registers (CSRs), which are defined in the Zicsr extension, and all other privileged instructions. The privileged architecture requires the Zicsr extension; which other privileged instructions are required depends on the privileged-architecture feature set.

In addition to the unprivileged state described in Volume I of this manual, an implementation may contain additional CSRs, accessible by some subset of the privilege levels using the CSR instructions described in Volume I. In this chapter, we map out the CSR address space. The following chapters describe the function of each of the CSRs according to privilege level, as well as the other privileged instructions which are generally closely associated with a particular privilege level. Note that although CSRs and instructions are associated with one privilege level, they are also accessible at all higher privilege levels.

Standard CSRs do not have side effects on reads but may have side effects on writes.

2.1. CSR Address Mapping Conventions

The standard RISC-V ISA sets aside a 12-bit encoding space (csr[11:0]) for up to 4,096 CSRs. By convention, the upper 4 bits of the CSR address (csr[11:8]) are used to encode the read and write accessibility of the CSRs according to privilege level as shown in Table 3. The top two bits (csr[11:10]) indicate whether the register is read/write (00,01, or 10) or read-only (11). The next two bits (csr[9:8]) encode the lowest privilege level that can access the CSR.

The CSR address convention uses the upper bits of the CSR address to encode default access privileges. This simplifies error checking in the hardware and provides a larger CSR space, but does constrain the mapping of CSRs into the address space. Implementations might allow a more-privileged level to trap otherwise permitted CSR accesses by a less-privileged level to allow these accesses to be intercepted. This change should be transparent to the less-privileged software.

Attempts to access a non-existent CSR raise an illegal-instruction exception. Attempts to access a CSR without appropriate privilege level or to write a read-only register also raise illegal-instruction exceptions. A read/write register might also contain some bits that are read-only, in which case writes to the read-only bits are ignored.

Table 3 also indicates the convention to allocate CSR addresses between standard and custom uses. The CSR addresses designated for custom uses will not be redefined by future standard extensions.

Machine-mode standard read-write CSRs 0x7A0-0x7BF are reserved for use by the debug system. Of these CSRs, 0x7A0-0x7AF are accessible to machine mode, whereas 0x7B0-0x7BF are only visible to debug mode. Implementations should raise illegal-instruction exceptions on machine-mode access to the latter set of registers.

Effective virtualization requires that as many instructions run natively as possible inside a virtualized environment, while any privileged accesses trap to the virtual machine monitor. (Goldberg, 1974) CSRs that are read-only at some lower privilege level are shadowed into separate CSR addresses if they are made read-write at a higher privilege level. This avoids trapping permitted lower-privilege accesses while still causing traps on illegal accesses. Currently, the counters are the only shadowed CSRs.

2.2. CSR Listing

Table 4-Table 8 list the CSRs that have currently been allocated CSR addresses. The timers, counters, and floating-point CSRs are standard unprivileged CSRs. The other registers are used by privileged code, as described in the following chapters. Note that not all registers are required on all implementations.

Table 3. Allocation of RISC-V CSR address ranges.

CSR Address			Hex	Use and Accessibility
[11:10]	[9:8]	[7:4]
Unprivileged and User-Level CSRs
00 01 10 11 11 11	00 00 00 00 00 00	XXXX XXXX XXXX 0XXX 10XX 11XX	0x000-0x0FF 0x400-0x4FF 0x800-0x8FF 0xC00-0xC7F 0xC80-0xCBF 0xCC0-0xCFF	Standard read/write Standard read/write Custom read/write Standard read-only Standard read-only Custom read-only
Supervisor-Level CSRs
00 01 01 01 10 10 10 11 11 11	01 01 01 01 01 01 01 01 01 01	XXXX 0XXX 10XX 11XX 0XXX 10XX 11XX 0XXX 10XX 11XX	0x100-0x1FF 0x500-0x57F 0x580-0x5BF 0x5C0-0x5FF 0x900-0x97F 0x980-0x9BF 0x9C0-0x9FF 0xD00-0xD7F 0xD80-0xDBF 0xDC0-0xDFF	Standard read/write Standard read/write Standard read/write Custom read/write Standard read/write Standard read/write Custom read/write Standard read-only Standard read-only Custom read-only
Hypervisor and VS CSRs
00 01 01 01 10 10 10 11 11 11	10 10 10 10 10 10 10 10 10 10	XXXX 0XXX 10XX 11XX 0XXX 10XX 11XX 0XXX 10XX 11XX	0x200-0x2FF 0x600-0x67F 0x680-0x6BF 0x6C0-0x6FF 0xA00-0xA7F 0xA80-0xABF 0xAC0-0xAFF 0xE00-0xE7F 0xE80-0xEBF 0xEC0-0xEFF	Standard read/write Standard read/write Standard read/write Custom read/write Standard read/write Standard read/write Custom read/write Standard read-only Standard read-only Custom read-only
Machine-Level CSRs
00 01 01 01 01 01 10 10 10 11 11 11	11 11 11 11 11 11 11 11 11 11 11 11	XXXX 0XXX 100X 1010 1011 11XX 0XXX 10XX 11XX 0XXX 10XX 11XX	0x300-0x3FF 0x700-0x77F 0x780-0x79F 0x7A0-0x7AF 0x7B0-0x7BF 0x7C0-0x7FF 0xB00-0xB7F 0xB80-0xBBF 0xBC0-0xBFF 0xF00-0xF7F 0xF80-0xFBF 0xFC0-0xFFF	Standard read/write Standard read/write Standard read/write Standard read/write debug CSRs Debug-mode-only CSRs Custom read/write Standard read/write Standard read/write Custom read/write Standard read-only Standard read-only Custom read-only

Table 3. Allocation of RISC-V CSR address ranges. {#csrrwpriv}

Table 4. Currently allocated RISC-V unprivileged CSR addresses.

Number	Privilege	Name	Description
Unprivileged Floating-Point CSRs
0x001 0x002 0x003	URW URW URW	fflags frm fcsr	Floating-Point Accrued Exceptions. Floating-Point Dynamic Rounding Mode. Floating-Point Control and Status Register (frm +fflags).
Unprivileged Counter/Timers
0xC00 0xC01 0xC02 0xC03 0xC04   0xC1F 0xC80 0xC81 0xC82 0xC83 0xC84   0xC9F	URO URO URO URO URO   URO URO URO URO URO URO   URO	cycle time instret hpmcounter3 hpmcounter4 ⋮ hpmcounter31 cycleh timeh instreth hpmcounter3h hpmcounter4h ⋮ hpmcounter31h	Cycle counter for RDCYCLE instruction. Timer for RDTIME instruction. Instructions-retired counter for RDINSTRET instruction. Performance-monitoring counter. Performance-monitoring counter.   Performance-monitoring counter. Upper 32 bits of cycle, RV32 only. Upper 32 bits of time, RV32 only. Upper 32 bits of instret, RV32 only. Upper 32 bits of hpmcounter3, RV32 only. Upper 32 bits of hpmcounter4, RV32 only.   Upper 32 bits of hpmcounter31, RV32 only.

Table 4. Currently allocated RISC-V unprivileged CSR addresses. {#ucsrnames}

Table 5. Currently allocated RISC-V supervisor-level CSR addresses.

Number	Privilege	Name	Description
Supervisor Trap Setup
0x100 0x104 0x105 0x106	SRW SRW SRW SRW	sstatus sie stvec scounteren	Supervisor status register. Supervisor interrupt-enable register. Supervisor trap handler base address. Supervisor counter enable.
Supervisor Configuration
0x10A	SRW	senvcfg	Supervisor environment configuration register.
Supervisor Trap Handling
0x140 0x141 0x142 0x143 0x144	SRW SRW SRW SRW SRW	sscratch sepc scause stval sip	Scratch register for supervisor trap handlers. Supervisor exception program counter. Supervisor trap cause. Supervisor bad address or instruction. Supervisor interrupt pending.
Supervisor Protection and Translation
0x180	SRW	satp	Supervisor address translation and protection.
Debug/Trace Registers
0x5A8	SRW	scontext	Supervisor-mode context register.

Table 5. Currently allocated RISC-V supervisor-level CSR addresses. {#scsrnames}

Table 6. Currently allocated RISC-V hypervisor and VS CSR addresses.

Number	Privilege	Name	Description
Hypervisor Trap Setup
0x600 0x602 0x603 0x604 0x606 0x607	HRW HRW HRW HRW HRW HRW	hstatus hedeleg hideleg hie hcounteren hgeie	Hypervisor status register. Hypervisor exception delegation register. Hypervisor interrupt delegation register. Hypervisor interrupt-enable register. Hypervisor counter enable. Hypervisor guest external interrupt-enable register.
Hypervisor Trap Handling
0x643 0x644 0x645 0x64A 0xE12	HRW HRW HRW HRW HRO	htval hip hvip htinst hgeip	Hypervisor bad guest physical address. Hypervisor interrupt pending. Hypervisor virtual interrupt pending. Hypervisor trap instruction (transformed). Hypervisor guest external interrupt pending.
Hypervisor Configuration
0x60A 0x61A	HRW HRM	henvcfg henvcfgh	Hypervisor environment configuration register. Upper 32 bits of henvcfg, RV32 only.
Hypervisor Protection and Translation
0x680	HRW	hgatp	Hypervisor guest address translation and protection.
Debug/Trace Registers
0x6A8	HRW	hcontext	Hypervisor-mode context register.
Hypervisor Counter/Timer Virtualization Registers
0x605 0x615	HRW HRW	htimedelta htimedeltah	Delta for VS/VU-mode timer. Upper 32 bits of htimedelta, RV32 only.
Virtual Supervisor Registers
0x200 0x204 0x205 0x240 0x241 0x242 0x243 0x244 0x280	HRW HRW HRW HRW HRW HRW HRW HRW HRW	vsstatus vsie vstvec vsscratch vsepc vscause vstval vsip vsatp	Virtual supervisor status register. Virtual supervisor interrupt-enable register. Virtual supervisor trap handler base address. Virtual supervisor scratch register. Virtual supervisor exception program counter. Virtual supervisor trap cause. Virtual supervisor bad address or instruction. Virtual supervisor interrupt pending. Virtual supervisor address translation and protection.

Table 6. Currently allocated RISC-V hypervisor and VS CSR addresses. {#hcsrnames}

Table 7. Currently allocated RISC-V machine-level CSR addresses.

Number	Privilege	Name	Description
Machine Information Registers
0xF11 0xF12 0xF13 0xF14 0xF15	MRO MRO MRO MRO MRO	mvendorid marchid mimpid mhartid mconfigptr	Vendor ID. Architecture ID. Implementation ID. Hardware thread ID. Pointer to configuration data structure.
Machine Trap Setup
0x300 0x301 0x302 0x303 0x304 0x305 0x306 0x310	MRW MRW MRW MRW MRW MRW MRW MRW	mstatus misa medeleg mideleg mie mtvec mcounteren mstatush	Machine status register. ISA and extensions Machine exception delegation register. Machine interrupt delegation register. Machine interrupt-enable register. Machine trap-handler base address. Machine counter enable. Additional machine status register, RV32 only.
Machine Trap Handling
0x340 0x341 0x342 0x343 0x344 0x34A 0x34B	MRW MRW MRW MRW MRW MRW MRW	mscratch mepc mcause mtval mip mtinst mtval2	Scratch register for machine trap handlers. Machine exception program counter. Machine trap cause. Machine bad address or instruction. Machine interrupt pending. Machine trap instruction (transformed). Machine bad guest physical address.
Machine Configuration
0x30A 0x31A 0x747 0x757	MRW MRW MRW MRW	menvcfg menvcfgh mseccfg mseccfgh	Machine environment configuration register. Upper 32 bits of menvcfg, RV32 only. Machine security configuration register. Upper 32 bits of mseccfg, RV32 only.
Machine Memory Protection
0x3A0 0x3A1 0x3A2 0x3A3   0x3AE 0x3AF 0x3B0 0x3B1   0x3EF	MRW MRW MRW MRW   MRW MRW MRW MRW   MRW	pmpcfg0 pmpcfg1 pmpcfg2 pmpcfg3 ⋯ pmpcfg14 pmpcfg15 pmpaddr0 pmpaddr1 ⋯ pmpaddr63	Physical memory protection configuration. Physical memory protection configuration, RV32 only. Physical memory protection configuration. Physical memory protection configuration, RV32 only.   Physical memory protection configuration. Physical memory protection configuration, RV32 only. Physical memory protection address register. Physical memory protection address register.   Physical memory protection address register.

Table 7. Currently allocated RISC-V machine-level CSR addresses. {#mcsrnames0}

Table 8. Currently allocated RISC-V machine-level CSR addresses.

Number	Privilege	Name	Description
Machine Non-Maskable Interrupt Handling
0x740 0x741 0x742 0x744	MRW MRW MRW MRW	mnscratch mnepc mncause mnstatus	Resumable NMI scratch register. Resumable NMI program counter. Resumable NMI cause. Resumable NMI status.
Machine Counter/Timers
0xB00 0xB02 0xB03 0xB04   0xB1F 0xB80 0xB82 0xB83 0xB84   0xB9F	MRW MRW MRW MRW   MRW MRW MRW MRW MRW   MRW	mcycle minstret mhpmcounter3 mhpmcounter4 ⋮ mhpmcounter31 mcycleh minstreth mhpmcounter3h mhpmcounter4h ⋮ mhpmcounter31h	Machine cycle counter. Machine instructions-retired counter. Machine performance-monitoring counter. Machine performance-monitoring counter.   Machine performance-monitoring counter. Upper 32 bits of mcycle, RV32 only. Upper 32 bits of minstret, RV32 only. Upper 32 bits of mhpmcounter3, RV32 only. Upper 32 bits of mhpmcounter4, RV32 only.   Upper 32 bits of mhpmcounter31, RV32 only.
Machine Counter Setup
0x320 0x323 0x324   0x33F	MRW MRW MRW   MRW	mcountinhibit mhpmevent3 mhpmevent4 ⋮ mhpmevent31	Machine counter-inhibit register. Machine performance-monitoring event selector. Machine performance-monitoring event selector.   Machine performance-monitoring event selector.
Debug/Trace Registers (shared with Debug Mode)
0x7A0 0x7A1 0x7A2 0x7A3 0x7A8	MRW MRW MRW MRW MRW	tselect tdata1 tdata2 tdata3 mcontext	Debug/Trace trigger register select. First Debug/Trace trigger data register. Second Debug/Trace trigger data register. Third Debug/Trace trigger data register. Machine-mode context register.
Debug Mode Registers
0x7B0 0x7B1 0x7B2 0x7B3	DRW DRW DRW DRW	dcsr dpc dscratch0 dscratch1	Debug control and status register. Debug program counter. Debug scratch register 0. Debug scratch register 1.

Table 8. Currently allocated RISC-V machine-level CSR addresses. {#mcsrnames1}

2.3. CSR Field Specifications

The following definitions and abbreviations are used in specifying the behavior of fields within the CSRs.

2.3.1. Reserved Writes Preserve Values, Reads Ignore Values (WPRI)

Some whole read/write fields are reserved for future use. Software should ignore the values read from these fields, and should preserve the values held in these fields when writing values to other fields of the same register. For forward compatibility, implementations that do not furnish these fields must make them read-only zero. These fields are labeled WPRI in the register descriptions.

To simplify the software model, any backward-compatible future definition of previously reserved fields within a CSR must cope with the possibility that a non-atomic read/modify/write sequence is used to update other fields in the CSR. Alternatively, the original CSR definition must specify that subfields can only be updated atomically, which may require a two-instruction clear bit/set bit sequence in general that can be problematic if intermediate values are not legal.

2.3.2. Write/Read Only Legal Values (WLRL)

Some read/write CSR fields specify behavior for only a subset of possible bit encodings, with other bit encodings reserved. Software should not write anything other than legal values to such a field, and should not assume a read will return a legal value unless the last write was of a legal value, or the register has not been written since another operation (e.g., reset) set the register to a legal value. These fields are labeled WLRL in the register descriptions.

Hardware implementations need only implement enough state bits to differentiate between the supported values, but must always return the complete specified bit-encoding of any supported value when read.

Implementations are permitted but not required to raise an illegal-instruction exception if an instruction attempts to write a non-supported value to a WLRL field. Implementations can return arbitrary bit patterns on the read of a WLRL field when the last write was of an illegal value, but the value returned should deterministically depend on the illegal written value and the value of the field prior to the write.

2.3.3. Write Any Values, Reads Legal Values (WARL)

Some read/write CSR fields are only defined for a subset of bit encodings, but allow any value to be written while guaranteeing to return a legal value whenever read. Assuming that writing the CSR has no other side effects, the range of supported values can be determined by attempting to write a desired setting then reading to see if the value was retained. These fields are labeled WARL in the register descriptions.

Implementations will not raise an exception on writes of unsupported values to a WARL field. Implementations can return any legal value on the read of a WARL field when the last write was of an illegal value, but the legal value returned should deterministically depend on the illegal written value and the architectural state of the hart.

2.4. CSR Field Modulation

If a write to one CSR changes the set of legal values allowed for a field of a second CSR, then unless specified otherwise, the second CSR’s field immediately gets an UNSPECIFIED value from among its new legal values. This is true even if the field’s value before the write remains legal after the write; the value of the field may be changed in consequence of the write to the controlling CSR.

As a special case of this rule, the value written to one CSR may control whether a field of a second CSR is writable (with multiple legal values) or is read-only. When a write to the controlling CSR causes the second CSR’s field to change from previously read-only to now writable, that field immediately gets an UNSPECIFIED but legal value, unless specified otherwise. Some CSR fields are, when writable, defined as aliases of other CSR fields. Let x be such a CSR field, and let y be the CSR field it aliases when writable. If a write to a controlling CSR causes field x to change from previously read-only to now writable, the new value of x is not UNSPECIFIED but instead immediately reflects the existing value of its alias y, as required.

A change to the value of a CSR for this reason is not a write to the affected CSR and thus does not trigger any side effects specified for that CSR.

2.5. Implicit Reads of CSRs

Implementations sometimes perform implicit reads of CSRs. (For example, all S-mode instruction fetches implicitly read the satp CSR.) Unless otherwise specified, the value returned by an implicit read of a CSR is the same value that would have been returned by an explicit read of the CSR, using a CSR-access instruction in a sufficient privilege mode.

2.6. CSR Width Modulation

If the width of a CSR is changed (for example, by changing MXLEN or UXLEN, as described in Section 3.1.6.2), the values of the writable fields and bits of the new-width CSR are, unless specified otherwise, determined from the previous-width CSR as though by this algorithm:

1. The value of the previous-width CSR is copied to a temporary register of the same width.

2. For the read-only bits of the previous-width CSR, the bits at the same positions in the temporary register are set to zeros.

3. The width of the temporary register is changed to the new width. If the new width W is narrower than the previous width, the least-significant W bits of the temporary register are retained and the more-significant bits are discarded. If the new width is wider than the previous width, the temporary register is zero-extended to the wider width.

4. Each writable field of the new-width CSR takes the value of the bits at the same positions in the temporary register.

Changing the width of a CSR is not a read or write of the CSR and thus does not trigger any side effects.

2.7. Explicit Accesses to CSRs Wider than XLEN

If a standard CSR is wider than XLEN bits, then an explicit read of the CSR returns the register’s least-significant XLEN bits, and an explicit write to the CSR modifies only the register’s least-significant XLEN bits, leaving the upper bits unchanged.

Some standard CSRs, such as the counter CSRs of extension Zicntr, are always 64 bits, even when XLEN=32 (RV32). For each such 64-bit CSR (for example, counter time), a corresponding 32-bit high-half CSR is usually defined with the same name but with the letter ‘h’ appended at the end (timeh). The high-half CSR aliases bits 63:32 of its namesake 64-bit CSR, thus providing a way for RV32 software to read and modify the otherwise-unreachable 32 bits.

Standard high-half CSRs are accessible only when the base RISC-V instruction set is RV32 (XLEN=32). For RV64 (when XLEN=64), the addresses of all standard high-half CSRs are reserved, so an attempt to access a high-half CSR typically raises an illegal-instruction exception.

3. Machine-Level ISA, Version 1.12

This chapter describes the machine-level operations available in machine-mode (M-mode), which is the highest privilege mode in a RISC-V system. M-mode is used for low-level access to a hardware platform and is the first mode entered at reset. M-mode can also be used to implement features that are too difficult or expensive to implement in hardware directly. The RISC-V machine-level ISA contains a common core that is extended depending on which other privilege levels are supported and other details of the hardware implementation.

3.1. Machine-Level CSRs

In addition to the machine-level CSRs described in this section, M-mode code can access all CSRs at lower privilege levels.

3.1.1. Machine ISA Register misa

The misa CSR is a WARL read-write register reporting the ISA supported by the hart. This register must be readable in any implementation, but a value of zero can be returned to indicate the misa register has not been implemented, requiring that CPU capabilities be determined through a separate non-standard mechanism.

Figure 2. Machine ISA register (misa)

The MXL (Machine XLEN) field encodes the native base integer ISA width as shown in Table 9. The MXL field may be writable in implementations that support multiple base ISAs. The effective XLEN in M-mode, MXLEN, is given by the setting of MXL, or has a fixed value if misa is zero. The MXL field is always set to the widest supported ISA variant at reset.

Table 9. Encoding of MXL field in misa

MXL	XLEN
1 2 3	32 64 128

Table 9. Encoding of MXL field in misa {#misabase}

The misa CSR is MXLEN bits wide. If the value read from misa is nonzero, field MXL of that value always denotes the current MXLEN. If a write to misa causes MXLEN to change, the position of MXL moves to the most-significant two bits of misa at the new width.

The base width can be quickly ascertained using branches on the sign of the returned misa value, and possibly a shift left by one and a second branch on the sign. These checks can be written in assembly code without knowing the register width (XLEN) of the machine. The base width is given by XLEN=2MXL+4. The base width can also be found if misa is zero, by placing the immediate 4 in a register then shifting the register left by 31 bits at a time. If zero after one shift, then the machine is RV32. If zero after two shifts, then the machine is RV64, else RV128.

The Extensions field encodes the presence of the standard extensions, with a single bit per letter of the alphabet (bit 0 encodes presence of extension "A" , bit 1 encodes presence of extension "B", through to bit 25 which encodes "Z"). The "I" bit will be set for RV32I, RV64I, and RV128I base ISAs, and the "E" bit will be set for RV32E and RV64E. The Extensions field is a WARL field that can contain writable bits where the implementation allows the supported ISA to be modified. At reset, the Extensions field shall contain the maximal set of supported extensions, and "I" shall be selected over "E" if both are available.

When a standard extension is disabled by clearing its bit in misa, the instructions and CSRs defined or modified by the extension revert to their defined or reserved behaviors as if the extension is not implemented.

For a given RISC-V execution environment, an instruction, extension, or other feature of the RISC-V ISA is ordinarily judged to be implemented or not by the observable execution behavior in that environment. For example, the F extension is said to be implemented for an execution environment if and only if the instructions that the RISC-V Unprivileged ISA defines for F execute as specified. With this definition of implemented, disabling an extension by clearing its bit in misa results in the extension being considered not implemented in M-mode. For example, setting misa.F=0 results in the F extension being not implemented for M-mode, because the F extension’s instructions will not act as the Unprivileged ISA requires but may instead raise an illegal-instruction exception. Defining the term implemented based strictly on the observable behavior might conflict with other common understandings of the same word. In particular, although common usage may allow for the combination "implemented but disabled," in this document it is considered a contradiction of terms, because disabled implies execution will not behave as required for the feature to be considered implemented. In the same vein, "implemented and enabled" is redundant here; "implemented" suffices.

Table 10. Encoding of Extensions field in misa. All bits that are reserved for future use must return zero when read.

Bit	Character	Description
0 1 2 3 4 5 6 7 8 9 10 11 12 13 14 15 16 17 18 19 20 21 22 23 24 25	A B C D E F G H I J K L M N O P Q R S T U V W X Y Z	Atomic extension Reserved Compressed extension Double-precision floating-point extension RV32E/64E base ISA Single-precision floating-point extension Reserved Hypervisor extension RV32I/64I/128I base ISA Reserved Reserved Reserved Integer Multiply/Divide extension Tentatively reserved for User-Level Interrupts extension Reserved Tentatively reserved for Packed-SIMD extension Quad-precision floating-point extension Reserved Supervisor mode implemented Reserved User mode implemented "V" Vector extension implemented Reserved Non-standard extensions present Reserved Reserved

Table 10. Encoding of Extensions field in misa. All bits that are reserved for future use must return zero when read.

The design of the RV128I base ISA is not yet complete, and while much of the remainder of this specification is expected to apply to RV128, this version of the document focuses only on RV32 and RV64.

The "U" and "S" bits will be set if there is support for user and supervisor modes respectively.

The "X" bit will be set if there are any non-standard extensions.

The misa CSR exposes a rudimentary catalog of CPU features to machine-mode code. More extensive information can be obtained in machine mode by probing other machine registers, and examining other ROM storage in the system as part of the boot process. We require that lower privilege levels execute environment calls instead of reading CPU registers to determine features available at each privilege level. This enables virtualization layers to alter the ISA observed at any level, and supports a much richer command interface without burdening hardware designs.

The "E" bit is read-only. Unless misa is all read-only zero, the "E" bit always reads as the complement of the "I" bit. If an execution environment supports both RV32E and RV32I, software can select RV32E by clearing the "I" bit.

If an ISA feature x depends on an ISA feature y, then attempting to enable feature x but disable feature y results in both features being disabled. For example, setting "F"=0 and "D"=1 results in both "F" and "D" being cleared.

An implementation may impose additional constraints on the collective setting of two or more misa fields, in which case they function collectively as a single WARL field. An attempt to write an unsupported combination causes those bits to be set to some supported combination.

Writing misa may increase IALIGN, e.g., by disabling the "C" extension. If an instruction that would write misa increases IALIGN, and the subsequent instruction’s address is not IALIGN-bit aligned, the write to misa is suppressed, leaving misa unchanged.

When software enables an extension that was previously disabled, then all state uniquely associated with that extension is UNSPECIFIED, unless otherwise specified by that extension.

3.1.2. Machine Vendor ID Register mvendorid

The mvendorid CSR is a 32-bit read-only register providing the JEDEC manufacturer ID of the provider of the core. This register must be readable in any implementation, but a value of 0 can be returned to indicate the field is not implemented or that this is a non-commercial implementation.

Figure 3. Vendor ID register (mvendorid)

JEDEC manufacturer IDs are ordinarily encoded as a sequence of one-byte continuation codes 0x7f, terminated by a one-byte ID not equal to 0x7f, with an odd parity bit in the most-significant bit of each byte. mvendorid encodes the number of one-byte continuation codes in the Bank field, and encodes the final byte in the Offset field, discarding the parity bit. For example, the JEDEC manufacturer ID 0x7f 0x7f 0x7f 0x7f 0x7f 0x7f 0x7f 0x7f 0x7f 0x7f 0x7f 0x7f 0x8a (twelve continuation codes followed by 0x8a) would be encoded in the mvendorid CSR as 0x60a.

In JEDEC’s parlance, the bank number is one greater than the number of continuation codes; hence, the mvendorid Bank field encodes a value that is one less than the JEDEC bank number. Previously the vendor ID was to be a number allocated by RISC-V International, but this duplicates the work of JEDEC in maintaining a manufacturer ID standard. At time of writing, registering a manufacturer ID with JEDEC has a one-time cost of $500.

3.1.3. Machine Architecture ID Register marchid

The marchid CSR is an MXLEN-bit read-only register encoding the base microarchitecture of the hart. This register must be readable in any implementation, but a value of 0 can be returned to indicate the field is not implemented. The combination of mvendorid and marchid should uniquely identify the type of hart microarchitecture that is implemented.

Figure 4. Machine Architecture ID register (marchid)

Open-source project architecture IDs are allocated globally by RISC-V International, and have non-zero architecture IDs with a zero most-significant-bit (MSB). Commercial architecture IDs are allocated by each commercial vendor independently, but must have the MSB set and cannot contain zero in the remaining MXLEN-1 bits.

The intent is for the architecture ID to represent the microarchitecture associated with the repo around which development occurs rather than a particular organization. Commercial fabrications of open-source designs should (and might be required by the license to) retain the original architecture ID. This will aid in reducing fragmentation and tool support costs, as well as provide attribution. Open-source architecture IDs are administered by RISC-V International and should only be allocated to released, functioning open-source projects. Commercial architecture IDs can be managed independently by any registered vendor but are required to have IDs disjoint from the open-source architecture IDs (MSB set) to prevent collisions if a vendor wishes to use both closed-source and open-source microarchitectures. The convention adopted within the following Implementation field can be used to segregate branches of the same architecture design, including by organization. The misa register also helps distinguish different variants of a design.

3.1.4. Machine Implementation ID Register mimpid

The mimpid CSR provides a unique encoding of the version of the processor implementation. This register must be readable in any implementation, but a value of 0 can be returned to indicate that the field is not implemented. The Implementation value should reflect the design of the RISC-V processor itself and not any surrounding system.

Figure 5. Machine Implementation ID register (mimpid)

The format of this field is left to the provider of the architecture source code, but will often be printed by standard tools as a hexadecimal string without any leading or trailing zeros, so the Implementation value can be left-justified (i.e., filled in from most-significant nibble down) with subfields aligned on nibble boundaries to ease human readability.

3.1.5. Hart ID Register mhartid

The mhartid CSR is an MXLEN-bit read-only register containing the integer ID of the hardware thread running the code. This register must be readable in any implementation. Hart IDs might not necessarily be numbered contiguously in a multiprocessor system, but at least one hart must have a hart ID of zero. Hart IDs must be unique within the execution environment.

Figure 6. Hart ID register (mhartid)

In certain cases, we must ensure exactly one hart runs some code (e.g., at reset), and so require one hart to have a known hart ID of zero. For efficiency, system implementers should aim to reduce the magnitude of the largest hart ID used in a system.

3.1.6. Machine Status Registers (mstatus and mstatush)

The mstatus register is an MXLEN-bit read/write register formatted as shown in Figure 7 for RV32 and Figure 8 for RV64. The mstatus register keeps track of and controls the hart’s current operating state. A restricted view of mstatus appears as the sstatus register in the S-level ISA.

Figure 7. Machine-mode status register (mstatus) for RV32

Figure 8. Machine-mode status register (mstatus) for RV64

For RV32 only, mstatush is a 32-bit read/write register formatted as shown in Figure 9. Bits 30:4 of mstatush generally contain the same fields found in bits 62:36 of mstatus for RV64. Fields SD, SXL, and UXL do not exist in mstatush.

Figure 9. Additional machine-mode status register (mstatush) for RV32.

3.1.6.1. Privilege and Global Interrupt-Enable Stack in mstatus register

Global interrupt-enable bits, MIE and SIE, are provided for M-mode and S-mode respectively. These bits are primarily used to guarantee atomicity with respect to interrupt handlers in the current privilege mode.

The global xIE bits are located in the low-order bits of mstatus, allowing them to be atomically set or cleared with a single CSR instruction.

When a hart is executing in privilege mode x, interrupts are globally enabled when xIE=1 and globally disabled when x IE=0. Interrupts for lower-privilege modes, w<x, are always globally disabled regardless of the setting of any global wIE bit for the lower-privilege mode. Interrupts for higher-privilege modes, y>x, are always globally enabled regardless of the setting of the global yIE bit for the higher-privilege mode. Higher-privilege-level code can use separate per-interrupt enable bits to disable selected higher-privilege-mode interrupts before ceding control to a lower-privilege mode.

A higher-privilege mode y could disable all of its interrupts before ceding control to a lower-privilege mode but this would be unusual as it would leave only a synchronous trap, non-maskable interrupt, or reset as means to regain control of the hart.

To support nested traps, each privilege mode x that can respond to interrupts has a two-level stack of interrupt-enable bits and privilege modes. xPIE holds the value of the interrupt-enable bit active prior to the trap, and xPP holds the previous privilege mode. The xPP fields can only hold privilege modes up to x, so MPP is two bits wide and SPP is one bit wide. When a trap is taken from privilege mode y into privilege mode x, xPIE is set to the value of xIE; xIE is set to 0; and xPP is set to y.

For lower privilege modes, any trap (synchronous or asynchronous) is usually taken at a higher privilege mode with interrupts disabled upon entry. The higher-level trap handler will either service the trap and return using the stacked information, or, if not returning immediately to the interrupted context, will save the privilege stack before re-enabling interrupts, so only one entry per stack is required.

An MRET or SRET instruction is used to return from a trap in M-mode or S-mode respectively. When executing an xRET instruction, supposing xPP holds the value y, xIE is set to xPIE; the privilege mode is changed to y; xPIE is set to 1; and xPP is set to the least-privileged supported mode (U if U-mode is implemented, else M). If y≠M, xRET also sets MPRV=0.

Setting xPP to the least-privileged supported mode on an xRET helps identify software bugs in the management of the two-level privilege-mode stack.

xPP fields are WARL fields that can hold only privilege mode x and any implemented privilege mode lower than x. If privilege mode x is not implemented, then xPP must be read-only 0.

M-mode software can determine whether a privilege mode is implemented by writing that mode to MPP then reading it back. If the machine provides only U and M modes, then only a single hardware storage bit is required to represent either 00 or 11 in MPP.

3.1.6.2. Base ISA Control in mstatus Register

For RV64 systems, the SXL and UXL fields are WARL fields that control the value of XLEN for S-mode and U-mode, respectively. The encoding of these fields is the same as the MXL field of misa, shown in Table 9. The effective XLEN in S-mode and U-mode are termed SXLEN and UXLEN, respectively.

When MXLEN=32, the SXL and UXL fields do not exist, and SXLEN=32 and UXLEN=32.

When MXLEN=64, if S-mode is not supported, then SXL is read-only zero. Otherwise, it is a WARL field that encodes the current value of SXLEN. In particular, an implementation may make SXL be a read-only field whose value always ensures that SXLEN=MXLEN.

When MXLEN=64, if U-mode is not supported, then UXL is read-only zero. Otherwise, it is a WARL field that encodes the current value of UXLEN. In particular, an implementation may make UXL be a read-only field whose value always ensures that UXLEN=MXLEN or UXLEN=SXLEN.

If S-mode is implemented, the set of legal values that the UXL field may assume excludes those that would cause UXLEN to be greater than SXLEN.

Whenever XLEN in any mode is set to a value less than the widest supported XLEN, all operations must ignore source operand register bits above the configured XLEN, and must sign-extend results to fill the entire widest supported XLEN in the destination register. Similarly, pc bits above XLEN are ignored, and when the pc is written, it is sign-extended to fill the widest supported XLEN.

We require that operations always fill the entire underlying hardware registers with defined values to avoid implementation-defined behavior. To reduce hardware complexity, the architecture imposes no checks that lower-privilege modes have XLEN settings less than or equal to the next-higher privilege mode. In practice, such settings would almost always be a software bug, but machine operation is well-defined even in this case.

If MXLEN is changed from 32 to a wider width, each of mstatus fields SXL and UXL, if not restricted to a single value, gets the value corresponding to the widest supported width not wider than the new MXLEN.

3.1.6.3. Memory Privilege in mstatus Register

The MPRV (Modify PRiVilege) bit modifies the effective privilege mode, i.e., the privilege level at which loads and stores execute. When MPRV=0, loads and stores behave as normal, using the translation and protection mechanisms of the current privilege mode. When MPRV=1, load and store memory addresses are translated and protected, and endianness is applied, as though the current privilege mode were set to MPP. Instruction address-translation and protection are unaffected by the setting of MPRV. MPRV is read-only 0 if U-mode is not supported.

An MRET or SRET instruction that changes the privilege mode to a mode less privileged than M also sets MPRV=0.

The MXR (Make eXecutable Readable) bit modifies the privilege with which loads access virtual memory. When MXR=0, only loads from pages marked readable (R=1 in Figure 57) will succeed. When MXR=1, loads from pages marked either readable or executable (R=1 or X=1) will succeed. MXR has no effect when page-based virtual memory is not in effect. MXR is read-only 0 if S-mode is not supported.

The MPRV and MXR mechanisms were conceived to improve the efficiency of M-mode routines that emulate missing hardware features, e.g., misaligned loads and stores. MPRV obviates the need to perform address translation in software. MXR allows instruction words to be loaded from pages marked execute-only. The current privilege mode and the privilege mode specified by MPP might have different XLEN settings. When MPRV=1, load and store memory addresses are treated as though the current XLEN were set to MPP’s XLEN, following the rules in Section 3.1.6.2.

The SUM (permit Supervisor User Memory access) bit modifies the privilege with which S-mode loads and stores access virtual memory. When SUM=0, S-mode memory accesses to pages that are accessible by U-mode (U=1 in Figure 57) will fault. When SUM=1, these accesses are permitted. SUM has no effect when page-based virtual memory is not in effect. Note that, while SUM is ordinarily ignored when not executing in S-mode, it is in effect when MPRV=1 and MPP=S. SUM is read-only 0 if S-mode is not supported or if satp.MODE is read-only 0.

The MXR and SUM mechanisms only affect the interpretation of permissions encoded in page-table entries. In particular, they have no impact on whether access-fault exceptions are raised due to PMAs or PMP.

3.1.6.4. Endianness Control in mstatus and mstatush Registers

The MBE, SBE, and UBE bits in mstatus and mstatush are WARL fields that control the endianness of memory accesses other than instruction fetches. Instruction fetches are always little-endian.

MBE controls whether non-instruction-fetch memory accesses made from M-mode (assuming mstatus.MPRV=0) are little-endian (MBE=0) or big-endian (MBE=1).

If S-mode is not supported, SBE is read-only 0. Otherwise, SBE controls whether explicit load and store memory accesses made from S-mode are little-endian (SBE=0) or big-endian (SBE=1).

If U-mode is not supported, UBE is read-only 0. Otherwise, UBE controls whether explicit load and store memory accesses made from U-mode are little-endian (UBE=0) or big-endian (UBE=1).

For implicit accesses to supervisor-level memory management data structures, such as page tables, endianness is always controlled by SBE. Since changing SBE alters the implementation’s interpretation of these data structures, if any such data structures remain in use across a change to SBE, M-mode software must follow such a change to SBE by executing an SFENCE.VMA instruction with rs1=x0 and rs2=x0.

Only in contrived scenarios will a given memory-management data structure be interpreted as both little-endian and big-endian. In practice, SBE will only be changed at runtime on world switches, in which case neither the old nor new memory-management data structure will be reinterpreted in a different endianness. In this case, no additional SFENCE.VMA is necessary, beyond what would ordinarily be required for a world switch.

If S-mode is supported, an implementation may make SBE be a read-only copy of MBE. If U-mode is supported, an implementation may make UBE be a read-only copy of either MBE or SBE.

An implementation supports only little-endian memory accesses if fields MBE, SBE, and UBE are all read-only 0. An implementation supports only big-endian memory accesses (aside from instruction fetches) if MBE is read-only 1 and SBE and UBE are each read-only 1 when S-mode and U-mode are supported. Volume I defines a hart’s address space as a circular sequence of 2XLEN bytes at consecutive addresses. The correspondence between addresses and byte locations is fixed and not affected by any endianness mode. Rather, the applicable endianness mode determines the order of mapping between memory bytes and a multibyte quantity (halfword, word, etc.). Standard RISC-V ABIs are expected to be purely little-endian-only or big-endian-only, with no accommodation for mixing endianness. Nevertheless, endianness control has been defined so as to permit, for instance, an OS of one endianness to execute user-mode programs of the opposite endianness. Consideration has been given also to the possibility of non-standard usages whereby software flips the endianness of memory accesses as needed. RISC-V instructions are uniformly little-endian to decouple instruction encoding from the current endianness settings, for the benefit of both hardware and software. Otherwise, for instance, a RISC-V assembler or disassembler would always need to know the intended active endianness, despite that the endianness mode might change dynamically during execution. In contrast, by giving instructions a fixed endianness, it is sometimes possible for carefully written software to be endianness-agnostic even in binary form, much like position-independent code. The choice to have instructions be only little-endian does have consequences, however, for RISC-V software that encodes or decodes machine instructions. In big-endian mode, such software must account for the fact that explicit loads and stores have endianness opposite that of instructions, for example by swapping byte order after loads and before stores.

3.1.6.5. Virtualization Support in mstatus Register

The TVM (Trap Virtual Memory) bit is a WARL field that supports intercepting supervisor virtual-memory management operations. When TVM=1, attempts to read or write the satp CSR or execute an SFENCE.VMA or SINVAL.VMA instruction while executing in S-mode will raise an illegal-instruction exception. When TVM=0, these operations are permitted in S-mode. TVM is read-only 0 when S-mode is not supported.

The TVM mechanism improves virtualization efficiency by permitting guest operating systems to execute in S-mode, rather than classically virtualizing them in U-mode. This approach obviates the need to trap accesses to most S-mode CSRs. Trapping satp accesses and the SFENCE.VMA and SINVAL.VMA instructions provides the hooks necessary to lazily populate shadow page tables.

The TW (Timeout Wait) bit is a WARL field that supports intercepting the WFI instruction (see Section 3.3.3). When TW=0, the WFI instruction may execute in lower privilege modes when not prevented for some other reason. When TW=1, then if WFI is executed in any less-privileged mode, and it does not complete within an implementation-specific, bounded time limit, the WFI instruction causes an illegal-instruction exception. An implementation may have WFI always raise an illegal-instruction exception in less-privileged modes when TW=1, even if there are pending globally-disabled interrupts when the instruction is executed. TW is read-only 0 when there are no modes less privileged than M.

Trapping the WFI instruction can trigger a world switch to another guest OS, rather than wastefully idling in the current guest.

When S-mode is implemented, then executing WFI in U-mode causes an illegal-instruction exception, unless it completes within an implementation-specific, bounded time limit. A future revision of this specification might add a feature that allows S-mode to selectively permit WFI in U-mode. Such a feature would only be active when TW=0.

The TSR (Trap SRET) bit is a WARL field that supports intercepting the supervisor exception return instruction, SRET. When TSR=1, attempts to execute SRET while executing in S-mode will raise an illegal-instruction exception. When TSR=0, this operation is permitted in S-mode. TSR is read-only 0 when S-mode is not supported.

Trapping SRET is necessary to emulate the hypervisor extension (see Chapter 9) on implementations that do not provide it.

3.1.6.6. Extension Context Status in mstatus Register

Supporting substantial extensions is one of the primary goals of RISC-V, and hence we define a standard interface to allow unchanged privileged-mode code, particularly a supervisor-level OS, to support arbitrary user-mode state extensions.

To date, the V extension is the only standard extension that defines additional state beyond the floating-point CSR and data registers.

The FS[1:0] and VS[1:0] WARL fields and the XS[1:0] read-only field are used to reduce the cost of context save and restore by setting and tracking the current state of the floating-point unit and any other user-mode extensions respectively. The FS field encodes the status of the floating-point unit state, including the floating-point registers f0–f31 and the CSRs fcsr, frm, and fflags. The VS field encodes the status of the vector extension state, including the vector registers v0–v31 and the CSRs vcsr, vxrm, vxsat, vstart, vl, vtype, and vlenb. The XS field encodes the status of additional user-mode extensions and associated state. These fields can be checked by a context switch routine to quickly determine whether a state save or restore is required. If a save or restore is required, additional instructions and CSRs are typically required to effect and optimize the process.

The design anticipates that most context switches will not need to save/restore state in either or both of the floating-point unit or other extensions, so provides a fast check via the SD bit.

The FS, VS, and XS fields use the same status encoding as shown in Table 11, with the four possible status values being Off, Initial, Clean, and Dirty.

Table 11. Encoding of FS[1:0], VS[1:0], and XS[1:0] status fields

Status	FS and VS Meaning	XS Meaning
0 1 2 3	Off Initial Clean Dirty	All off None dirty or clean, some on None dirty, some clean Some dirty

Table 11. Encoding of FS[1:0], VS[1:0], and XS[1:0] status fields {#fsxsencoding}

If the F extension is implemented, the FS field shall not be read-only zero.

If neither the F extension nor S-mode is implemented, then FS is read-only zero. If S-mode is implemented but the F extension is not, FS may optionally be read-only zero.

Implementations with S-mode but without the F extension are permitted, but not required, to make the FS field be read-only zero. Some such implementations will choose not to have the FS field be read-only zero, so as to enable emulation of the F extension for both S-mode and U-mode via invisible traps into M-mode.

If the v registers are implemented, the VS field shall not be read-only zero.

If neither the v registers nor S-mode is implemented, then VS is read-only zero. If S-mode is implemented but the v registers are not, VS may optionally be read-only zero.

In systems without additional user extensions requiring new state, the XS field is read-only zero. Every additional extension with state provides a CSR field that encodes the equivalent of the XS states. The XS field represents a summary of all extensions' status as shown in Table 11.

The XS field effectively reports the maximum status value across all user-extension status fields, though individual extensions can use a different encoding than XS.

The SD bit is a read-only bit that summarizes whether either the FS, VS, or XS fields signal the presence of some dirty state that will require saving extended user context to memory. If FS, XS, and VS are all read-only zero, then SD is also always zero.

When an extension’s status is set to Off, any instruction that attempts to read or write the corresponding state will cause an illegal-instruction exception. When the status is Initial, the corresponding state should have an initial constant value. When the status is Clean, the corresponding state is potentially different from the initial value, but matches the last value stored on a context swap. When the status is Dirty, the corresponding state has potentially been modified since the last context save.

During a context save, the responsible privileged code need only write out the corresponding state if its status is Dirty, and can then reset the extension’s status to Clean. During a context restore, the context need only be loaded from memory if the status is Clean (it should never be Dirty at restore). If the status is Initial, the context must be set to an initial constant value on context restore to avoid a security hole, but this can be done without accessing memory. For example, the floating-point registers can all be initialized to the immediate value 0.

The FS and XS fields are read by the privileged code before saving the context. The FS field is set directly by privileged code when resuming a user context, while the XS field is set indirectly by writing to the status register of the individual extensions. The status fields will also be updated during execution of instructions, regardless of privilege mode.

Extensions to the user-mode ISA often include additional user-mode state, and this state can be considerably larger than the base integer registers. The extensions might only be used for some applications, or might only be needed for short phases within a single application. To improve performance, the user-mode extension can define additional instructions to allow user-mode software to return the unit to an initial state or even to turn off the unit.

For example, a coprocessor might require to be configured before use and can be "unconfigured" after use. The unconfigured state would be represented as the Initial state for context save. If the same application remains running between the unconfigure and the next configure (which would set status to Dirty), there is no need to actually reinitialize the state at the unconfigure instruction, as all state is local to the user process, i.e., the Initial state may only cause the coprocessor state to be initialized to a constant value at context restore, not at every unconfigure.

Executing a user-mode instruction to disable a unit and place it into the Off state will cause an illegal-instruction exception to be raised if any subsequent instruction tries to use the unit before it is turned back on. A user-mode instruction to turn a unit on must also ensure the unit’s state is properly initialized, as the unit might have been used by another context meantime.

Changing the setting of FS has no effect on the contents of the floating-point register state. In particular, setting FS=Off does not destroy the state, nor does setting FS=Initial clear the contents. Similarly, the setting of VS has no effect on the contents of the vector register state. Other extensions, however, might not preserve state when set to Off.

Implementations may choose to track the dirtiness of the floating-point register state imprecisely by reporting the state to be dirty even when it has not been modified. On some implementations, some instructions that do not mutate the floating-point state may cause the state to transition from Initial or Clean to Dirty. On other implementations, dirtiness might not be tracked at all, in which case the valid FS states are Off and Dirty, and an attempt to set FS to Initial or Clean causes it to be set to Dirty.

This definition of FS does not disallow setting FS to Dirty as a result of errant speculation. Some platforms may choose to disallow speculatively writing FS to close a potential side channel.

If an instruction explicitly or implicitly writes a floating-point register or the fcsr but does not alter its contents, and FS=Initial or FS=Clean, it is implementation-defined whether FS transitions to Dirty.

Implementations may choose to track the dirtiness of the vector register state in an analogous imprecise fashion, including possibly setting VS to Dirty when software attempts to set VS=Initial or VS=Clean. When VS=Initial or VS=Clean, it is implementation-defined whether an instruction that writes a vector register or vector CSR but does not alter its contents causes VS to transition to Dirty.

Table 12 shows all the possible state transitions for the FS, VS, or XS status bits. Note that the standard floating-point and vector extensions do not support user-mode unconfigure or disable/enable instructions.

Current State Action

Off

Initial

Clean

Dirty

At context save in privileged code
Save state? Next state	No Off	No Initial	No Clean	Yes Clean

At context restore in privileged code
Restore state? Next state	No Off	Yes, to initial Initial	Yes, from memory Clean	N/A N/A

Execute instruction to read state
Action? Next state	Exception Off	Execute Initial	Execute Clean	Execute Dirty

Execute instruction that possibly modifies state, including configuration
Action? Next state	Exception Off	Execute Dirty	Execute Dirty	Execute Dirty

Execute instruction to unconfigure unit
Action? Next state	Exception Off	Execute Initial	Execute Initial	Execute Initial

Execute instruction to disable unit
Action? Next state	Execute Off	Execute Off	Execute Off	Execute Off

Table 12. FS, FS, and XS state transitions.

Execute instruction to enable unit
Action? Next state	Execute Initial	Execute Initial	Execute Initial	Execute Initial

Table 12. FS, FS, and XS state transitions. {#fsxsstates}

Standard privileged instructions to initialize, save, and restore extension state are provided to insulate privileged code from details of the added extension state by treating the state as an opaque object.

Many coprocessor extensions are only used in limited contexts that allows software to safely unconfigure or even disable units when done. This reduces the context-switch overhead of large stateful coprocessors. We separate out floating-point state from other extension state, as when a floating-point unit is present the floating-point registers are part of the standard calling convention, and so user-mode software cannot know when it is safe to disable the floating-point unit.

The XS field provides a summary of all added extension state, but additional microarchitectural bits might be maintained in the extension to further reduce context save and restore overhead.

The SD bit is read-only and is set when either the FS, VS, or XS bits encode a Dirty state (i.e., SD=FS==11) OR (XS==11) OR (VS==11). This allows privileged code to quickly determine when no additional context save is required beyond the integer register set and pc.

The floating-point unit state is always initialized, saved, and restored using standard instructions (F, D, and/or Q), and privileged code must be aware of FLEN to determine the appropriate space to reserve for each f register.

Machine and Supervisor modes share a single copy of the FS, VS, and XS bits. Supervisor-level software normally uses the FS, VS, and XS bits directly to record the status with respect to the supervisor-level saved context. Machine-level software must be more conservative in saving and restoring the extension state in their corresponding version of the context.

In any reasonable use case, the number of context switches between user and supervisor level should far outweigh the number of context switches to other privilege levels. Note that coprocessors should not require their context to be saved and restored to service asynchronous interrupts, unless the interrupt results in a user-level context swap.

3.1.7. Machine Trap-Vector Base-Address Register (mtvec)

The mtvec register is an MXLEN-bit WARL read/write register that holds trap vector configuration, consisting of a vector base address (BASE) and a vector mode (MODE).

Figure 10. Encoding of mtvec MODE field.

The mtvec register must always be implemented, but can contain a read-only value. If mtvec is writable, the set of values the register may hold can vary by implementation. The value in the BASE field must always be aligned on a 4-byte boundary, and the MODE setting may impose additional alignment constraints on the value in the BASE field.

We allow for considerable flexibility in implementation of the trap vector base address. On the one hand, we do not wish to burden low-end implementations with a large number of state bits, but on the other hand, we wish to allow flexibility for larger systems.

Table 13. Encoding of mtvec MODE field.

Value	Name	Description
0 1 ≥2	Direct Vectored ---	All exceptions set pc to BASE. Asynchronous interrupts set pc to BASE+4×cause. Reserved

Table 13. Encoding of mtvec MODE field. {#mtvec-mode}

The encoding of the MODE field is shown in Table 13. When MODE=Direct, all traps into machine mode cause the pc to be set to the address in the BASE field. When MODE=Vectored, all synchronous exceptions into machine mode cause the pc to be set to the address in the BASE field, whereas interrupts cause the pc to be set to the address in the BASE field plus four times the interrupt cause number. For example, a machine-mode timer interrupt (see Table 14) causes the pc to be set to BASE+0x1c.

When vectored interrupts are enabled, interrupt cause 0, which corresponds to user-mode software interrupts, are vectored to the same location as synchronous exceptions. This ambiguity does not arise in practice, since user-mode software interrupts are either disabled or delegated to user mode.

An implementation may have different alignment constraints for different modes. In particular, MODE=Vectored may have stricter alignment constraints than MODE=Direct.

Allowing coarser alignments in Vectored mode enables vectoring to be implemented without a hardware adder circuit. Reset and NMI vector locations are given in a platform specification.

3.1.8. Machine Trap Delegation Registers (medeleg and mideleg)

By default, all traps at any privilege level are handled in machine mode, though a machine-mode handler can redirect traps back to the appropriate level with the MRET instruction (Section 3.3.2). To increase performance, implementations can provide individual read/write bits within medeleg and mideleg to indicate that certain exceptions and interrupts should be processed directly by a lower privilege level. The machine exception delegation register (medeleg) and machine interrupt delegation register (mideleg) are MXLEN-bit read/write registers.

In systems with S-mode, the medeleg and mideleg registers must exist, and setting a bit in medeleg or mideleg will delegate the corresponding trap, when occurring in S-mode or U-mode, to the S-mode trap handler. In systems without S-mode, the medeleg and mideleg registers should not exist.

In versions 1.9.1 and earlier , these registers existed but were hardwired to zero in M-mode only, or M/U without N systems. There is no reason to require they return zero in those cases, as the misa register indicates whether they exist.

When a trap is delegated to S-mode, the scause register is written with the trap cause; the sepc register is written with the virtual address of the instruction that took the trap; the stval register is written with an exception-specific datum; the SPP field of mstatus is written with the active privilege mode at the time of the trap; the SPIE field of mstatus is written with the value of the SIE field at the time of the trap; and the SIE field of mstatus is cleared. The mcause, mepc, and mtval registers and the MPP and MPIE fields of mstatus are not written.

An implementation can choose to subset the delegatable traps, with the supported delegatable bits found by writing one to every bit location, then reading back the value in medeleg or mideleg to see which bit positions hold a one.

An implementation shall not have any bits of medeleg be read-only one, i.e., any synchronous trap that can be delegated must support not being delegated. Similarly, an implementation shall not fix as read-only one any bits of mideleg corresponding to machine-level interrupts (but may do so for lower-level interrupts).

Version 1.11 and earlier prohibited having any bits of mideleg be read-only one. Platform standards may always add such restrictions.

Traps never transition from a more-privileged mode to a less-privileged mode. For example, if M-mode has delegated illegal-instruction exceptions to S-mode, and M-mode software later executes an illegal instruction, the trap is taken in M-mode, rather than being delegated to S-mode. By contrast, traps may be taken horizontally. Using the same example, if M-mode has delegated illegal-instruction exceptions to S-mode, and S-mode software later executes an illegal instruction, the trap is taken in S-mode.

Delegated interrupts result in the interrupt being masked at the delegator privilege level. For example, if the supervisor timer interrupt (STI) is delegated to S-mode by setting mideleg[5], STIs will not be taken when executing in M-mode. By contrast, if mideleg[5] is clear, STIs can be taken in any mode and regardless of current mode will transfer control to M-mode.

Figure 11. Machine Exception Delegation Register medeleg.

medeleg has a bit position allocated for every synchronous exception shown in Table 14, with the index of the bit position equal to the value returned in the mcause register (i.e., setting bit 8 allows user-mode environment calls to be delegated to a lower-privilege trap handler).

Figure 12. Machine Interrupt Delegation Register mideleg.

mideleg holds trap delegation bits for individual interrupts, with the layout of bits matching those in the mip register (i.e., STIP interrupt delegation control is located in bit 5).

For exceptions that cannot occur in less privileged modes, the corresponding medeleg bits should be read-only zero. In particular, medeleg[11] is read-only zero.

3.1.9. Machine Interrupt Registers (mip and mie)

The mip register is an MXLEN-bit read/write register containing information on pending interrupts, while mie is the corresponding MXLEN-bit read/write register containing interrupt enable bits. Interrupt cause number i (as reported in CSR mcause, Section 3.1.15) corresponds with bit i in both mip and mie. Bits 15:0 are allocated to standard interrupt causes only, while bits 16 and above are designated for platform or custom use.

Figure 13. Machine Interrupt-Pending Register (mip).

Figure 14. Machine Interrupt-Enable Register (mie)

An interrupt i will trap to M-mode (causing the privilege mode to change to M-mode) if all of the following are true: (a) either the current privilege mode is M and the MIE bit in the mstatus register is set, or the current privilege mode has less privilege than M-mode; (b) bit i is set in both mip and mie; and (c) if register mideleg exists, bit i is not set in mideleg.

These conditions for an interrupt trap to occur must be evaluated in a bounded amount of time from when an interrupt becomes, or ceases to be, pending in mip, and must also be evaluated immediately following the execution of an xRET instruction or an explicit write to a CSR on which these interrupt trap conditions expressly depend (including mip, mie, mstatus, and mideleg).

Interrupts to M-mode take priority over any interrupts to lower privilege modes.

Each individual bit in register mip may be writable or may be read-only. When bit i in mip is writable, a pending interrupt i can be cleared by writing 0 to this bit. If interrupt i can become pending but bit i in mip is read-only, the implementation must provide some other mechanism for clearing the pending interrupt.

A bit in mie must be writable if the corresponding interrupt can ever become pending. Bits of mie that are not writable must be read-only zero.

The standard portions (bits 15:0) of registers mip and mie are formatted as shown in Figure 15 and Figure 16 respectively.

Figure 15. Standard portion (bits 15:0) of mip.

Figure 16. Standard portion (bits 15:0) of mie.

The machine-level interrupt registers handle a few root interrupt sources which are assigned a fixed service priority for simplicity, while separate external interrupt controllers can implement a more complex prioritization scheme over a much larger set of interrupts that are then muxed into the machine-level interrupt sources. The non-maskable interrupt is not made visible via the mip register as its presence is implicitly known when executing the NMI trap handler.

Bits mip.MEIP and mie.MEIE are the interrupt-pending and interrupt-enable bits for machine-level external interrupts. MEIP is read-only in mip, and is set and cleared by a platform-specific interrupt controller.

Bits mip.MTIP and mie.MTIE are the interrupt-pending and interrupt-enable bits for machine timer interrupts. MTIP is read-only in mip, and is cleared by writing to the memory-mapped machine-mode timer compare register.

Bits mip.MSIP and mie.MSIE are the interrupt-pending and interrupt-enable bits for machine-level software interrupts. MSIP is read-only in mip, and is written by accesses to memory-mapped control registers, which are used by remote harts to provide machine-level interprocessor interrupts. A hart can write its own MSIP bit using the same memory-mapped control register. If a system has only one hart, or if a platform standard supports the delivery of machine-level interprocessor interrupts through external interrupts (MEI) instead, then mip.MSIP and mie.MSIE may both be read-only zeros.

If supervisor mode is not implemented, bits SEIP, STIP, and SSIP of mip and SEIE, STIE, and SSIE of mie are read-only zeros.

If supervisor mode is implemented, bits mip.SEIP and mie.SEIE are the interrupt-pending and interrupt-enable bits for supervisor-level external interrupts. SEIP is writable in mip, and may be written by M-mode software to indicate to S-mode that an external interrupt is pending. Additionally, the platform-level interrupt controller may generate supervisor-level external interrupts. Supervisor-level external interrupts are made pending based on the logical-OR of the software-writable SEIP bit and the signal from the external interrupt controller. When mip is read with a CSR instruction, the value of the SEIP bit returned in the rd destination register is the logical-OR of the software-writable bit and the interrupt signal from the interrupt controller, but the signal from the interrupt controller is not used to calculate the value written to SEIP. Only the software-writable SEIP bit participates in the read-modify-write sequence of a CSRRS or CSRRC instruction.

For example, if we name the software-writable SEIP bit B and the signal from the external interrupt controller E, then if csrrs t0, mip, t1 is executed, t0[9] is written with B || E, then B is written with B || t1[9]. If csrrw t0, mip, t1 is executed, then t0[9] is written with B || E, and B is simply written with t1[9]. In neither case does B depend upon E. The SEIP field behavior is designed to allow a higher privilege layer to mimic external interrupts cleanly, without losing any real external interrupts. The behavior of the CSR instructions is slightly modified from regular CSR accesses as a result.

If supervisor mode is implemented, bits mip.STIP and mie.STIE are the interrupt-pending and interrupt-enable bits for supervisor-level timer interrupts. STIP is writable in mip, and may be written by M-mode software to deliver timer interrupts to S-mode.

If supervisor mode is implemented, bits mip.SSIP and mie.SSIE are the interrupt-pending and interrupt-enable bits for supervisor-level software interrupts. SSIP is writable in mip and may also be set to 1 by a platform-specific interrupt controller.

Multiple simultaneous interrupts destined for M-mode are handled in the following decreasing priority order: MEI, MSI, MTI, SEI, SSI, STI.

The machine-level interrupt fixed-priority ordering rules were developed with the following rationale. Interrupts for higher privilege modes must be serviced before interrupts for lower privilege modes to support preemption. The platform-specific machine-level interrupt sources in bits 16 and above have platform-specific priority, but are typically chosen to have the highest service priority to support very fast local vectored interrupts. External interrupts are handled before internal (timer/software) interrupts as external interrupts are usually generated by devices that might require low interrupt service times. Software interrupts are handled before internal timer interrupts, because internal timer interrupts are usually intended for time slicing, where time precision is less important, whereas software interrupts are used for inter-processor messaging. Software interrupts can be avoided when high-precision timing is required, or high-precision timer interrupts can be routed via a different interrupt path. Software interrupts are located in the lowest four bits of mip as these are often written by software, and this position allows the use of a single CSR instruction with a five-bit immediate.

Restricted views of the mip and mie registers appear as the sip and sie registers for supervisor level. If an interrupt is delegated to S-mode by setting a bit in the mideleg register, it becomes visible in the sip register and is maskable using the sie register. Otherwise, the corresponding bits in sip and sie are read-only zero.

3.1.10. Hardware Performance Monitor

M-mode includes a basic hardware performance-monitoring facility. The mcycle CSR counts the number of clock cycles executed by the processor core on which the hart is running. The minstret CSR counts the number of instructions the hart has retired. The mcycle and minstret registers have 64-bit precision on all RV32 and RV64 systems.

The counter registers have an arbitrary value after the hart is reset, and can be written with a given value. Any CSR write takes effect after the writing instruction has otherwise completed. The mcycle CSR may be shared between harts on the same core, in which case writes to mcycle will be visible to those harts. The platform should provide a mechanism to indicate which harts share an mcycle CSR.

The hardware performance monitor includes 29 additional 64-bit event counters, mhpmcounter3-mhpmcounter31. The event selector CSRs, mhpmevent3-mhpmevent31, are MXLEN-bit WARL registers that control which event causes the corresponding counter to increment. The meaning of these events is defined by the platform, but event 0 is defined to mean "no event." All counters should be implemented, but a legal implementation is to make both the counter and its corresponding event selector be read-only 0.

Figure 17. Hardware performance monitor counters.

The mhpmcounters are WARL registers that support up to 64 bits of precision on RV32 and RV64.

A future revision of this specification will define a mechanism to generate an interrupt when a hardware performance monitor counter overflows.

When XLEN=32, reads of the mcycle, minstret, and mhpmcountern CSRs return bits 31-0 of the corresponding counter, and writes change only bits 31-0; reads of the mcycleh, minstreth, and mhpmcounternh CSRs return bits 63-32 of the corresponding counter, and writes change only bits 63-32.

Figure 18. Upper 32 bits of hardware performance monitor counters, RV32 only.

3.1.11. Machine Counter-Enable Register (mcounteren)

The counter-enable register mcounteren is a 32-bit register that controls the availability of the hardware performance-monitoring counters to the next-lower privileged mode.

Figure 19. Counter-enable register (mcounteren).

The settings in this register only control accessibility. The act of reading or writing this register does not affect the underlying counters, which continue to increment even when not accessible.

When the CY, TM, IR, or HPMn bit in the mcounteren register is clear, attempts to read the cycle, time, instret, or hpmcountern register while executing in S-mode or U-mode will cause an illegal-instruction exception. When one of these bits is set, access to the corresponding register is permitted in the next implemented privilege mode (S-mode if implemented, otherwise U-mode).

The counter-enable bits support two common use cases with minimal hardware. For systems that do not need high-performance timers and counters, machine-mode software can trap accesses and implement all features in software. For systems that need high-performance timers and counters but are not concerned with obfuscating the underlying hardware counters, the counters can be directly exposed to lower privilege modes.

The cycle, instret, and hpmcountern CSRs are read-only shadows of mcycle, minstret, and mhpmcounter n, respectively. The time CSR is a read-only shadow of the memory-mapped mtime register. Analogously, on RV32I the cycleh, instreth and hpmcounternh CSRs are read-only shadows of mcycleh, minstreth and mhpmcounternh, respectively. On RV32I the timeh CSR is a read-only shadow of the upper 32 bits of the memory-mapped mtime register, while time shadows only the lower 32 bits of mtime.

Implementations can convert reads of the time and timeh CSRs into loads to the memory-mapped mtime register, or emulate this functionality on behalf of less-privileged modes in M-mode software.

In systems with U-mode, the mcounteren must be implemented, but all fields are WARL and may be read-only zero, indicating reads to the corresponding counter will cause an illegal-instruction exception when executing in a less-privileged mode. In systems without U-mode, the mcounteren register should not exist.

3.1.12. Machine Counter-Inhibit CSR (mcountinhibit)

Figure 20. Counter-inhibit register mcountinhibit

The counter-inhibit register mcountinhibit is a 32-bit WARL register that controls which of the hardware performance-monitoring counters increment. The settings in this register only control whether the counters increment; their accessibility is not affected by the setting of this register.

When the CY, IR, or HPMn bit in the mcountinhibit register is clear, the cycle, instret, or hpmcountern register increments as usual. When the CY, IR, or HPM_n_ bit is set, the corresponding counter does not increment.

The mcycle CSR may be shared between harts on the same core, in which case the mcountinhibit.CY field is also shared between those harts, and so writes to mcountinhibit.CY will be visible to those harts.

If the mcountinhibit register is not implemented, the implementation behaves as though the register were set to zero.

When the cycle and instret counters are not needed, it is desirable to conditionally inhibit them to reduce energy consumption. Providing a single CSR to inhibit all counters also allows the counters to be atomically sampled. Because the time counter can be shared between multiple cores, it cannot be inhibited with the mcountinhibit mechanism.

3.1.13. Machine Scratch Register (mscratch)

The mscratch register is an MXLEN-bit read/write register dedicated for use by machine mode. Typically, it is used to hold a pointer to a machine-mode hart-local context space and swapped with a user register upon entry to an M-mode trap handler.

Figure 21. Machine-mode scratch register.

The MIPS ISA allocated two user registers (k0/k1) for use by the operating system. Although the MIPS scheme provides a fast and simple implementation, it also reduces available user registers, and does not scale to further privilege levels, or nested traps. It can also require both registers are cleared before returning to user level to avoid a potential security hole and to provide deterministic debugging behavior. The RISC-V user ISA was designed to support many possible privileged system environments and so we did not want to infect the user-level ISA with any OS-dependent features. The RISC-V CSR swap instructions can quickly save/restore values to the mscratch register. Unlike the MIPS design, the OS can rely on holding a value in the mscratch register while the user context is running.

3.1.14. Machine Exception Program Counter (mepc)

mepc is an MXLEN-bit read/write register formatted as shown in Figure 22. The low bit of mepc (mepc[0]) is always zero. On implementations that support only IALIGN=32, the two low bits (mepc[1:0]) are always zero.

If an implementation allows IALIGN to be either 16 or 32 (by changing CSR misa, for example), then, whenever IALIGN=32, bit mepc[1] is masked on reads so that it appears to be 0. This masking occurs also for the implicit read by the MRET instruction. Though masked, mepc[1] remains writable when IALIGN=32.

mepc is a WARL register that must be able to hold all valid virtual addresses. It need not be capable of holding all possible invalid addresses. Prior to writing mepc, implementations may convert an invalid address into some other invalid address that mepc is capable of holding.

When address translation is not in effect, virtual addresses and physical addresses are equal. Hence, the set of addresses mepc must be able to represent includes the set of physical addresses that can be used as a valid pc or effective address.

When a trap is taken into M-mode, mepc is written with the virtual address of the instruction that was interrupted or that encountered the exception. Otherwise, mepc is never written by the implementation, though it may be explicitly written by software.

Figure 22. Machine exception program counter register.

3.1.15. Machine Cause Register (mcause)

The mcause register is an MXLEN-bit read-write register formatted as shown in Figure 23. When a trap is taken into M-mode, mcause is written with a code indicating the event that caused the trap. Otherwise, mcause is never written by the implementation, though it may be explicitly written by software.

The Interrupt bit in the mcause register is set if the trap was caused by an interrupt. The Exception Code field contains a code identifying the last exception or interrupt. Table 14 lists the possible machine-level exception codes. The Exception Code is a WLRL field, so is only guaranteed to hold supported exception codes.

Figure 23. Machine Cause register mcause.

Note that load and load-reserved instructions generate load exceptions, whereas store, store-conditional, and AMO instructions generate store/AMO exceptions.

Interrupts can be separated from other traps with a single branch on the sign of the mcause register value. A shift left can remove the interrupt bit and scale the exception codes to index into a trap vector table. We do not distinguish privileged instruction exceptions from illegal-instruction exceptions. This simplifies the architecture and also hides details of which higher-privilege instructions are supported by an implementation. The privilege level servicing the trap can implement a policy on whether these need to be distinguished, and if so, whether a given opcode should be treated as illegal or privileged.

If an instruction may raise multiple synchronous exceptions, the decreasing priority order of Table 15 indicates which exception is taken and reported in mcause. The priority of any custom synchronous exceptions is implementation-defined.

Table 14. Machine cause register (mcause) values after trap.

Interrupt	Exception Code	Description
1 1 1 1	0 1 2 3	Reserved Supervisor software interrupt Reserved Machine software interrupt
1 1 1 1	4 5 6 7	Reserved Supervisor timer interrupt Reserved Machine timer interrupt
1 1 1 1	8 9 10 11	Reserved Supervisor external interrupt Reserved Machine external interrupt
1 1	12-15 ≥16	Reserved Designated for platform use
0 0 0 0 0 0 0 0 0 0 0 0 0 0 0 0 0 0 0 0 0	0 1 2 3 4 5 6 7 8 9 10 11 12 13 14 15 16-23 24-31 32-47 48-63 ≥64	Instruction address misaligned Instruction access fault Illegal instruction Breakpoint Load address misaligned Load access fault Store/AMO address misaligned Store/AMO access fault Environment call from U-mode Environment call from S-mode Reserved Environment call from M-mode Instruction page fault Load page fault Reserved Store/AMO page fault Reserved Designated for custom use Reserved Designated for custom use Reserved

Table 14. Machine cause register (mcause) values after trap. {#mcauses}

Table 15. Synchronous exception priority in decreasing priority order.

Priority	Exc.Code	Description
Highest	3	Instruction address breakpoint
	12, 1	During instruction address translation: First encountered page fault or access fault
	1	With physical address for instruction: Instruction access fault
	2 0 8,9,11 3 3	Illegal instruction Instruction address misaligned Environment call Environment break Load/store/AMO address breakpoint
	4,6	Optionally: Load/store/AMO address misaligned
	13, 15, 5, 7	During address translation for an explicit memory access: First encountered page fault or access fault
	5,7	With physical address for an explicit memory access: Load/store/AMO access fault
Lowest	4,6	If not higher priority: Load/store/AMO address misaligned

Table 15. Synchronous exception priority in decreasing priority order. {#exception-priority}

When a virtual address is translated into a physical address, the address translation algorithm determines what specific exception may be raised.

Load/store/AMO address-misaligned exceptions may have either higher or lower priority than load/store/AMO page-fault and access-fault exceptions.

The relative priority of load/store/AMO address-misaligned and page-fault exceptions is implementation-defined to flexibly cater to two design points. Implementations that never support misaligned accesses can unconditionally raise the misaligned-address exception without performing address translation or protection checks. Implementations that support misaligned accesses only to some physical addresses must translate and check the address before determining whether the misaligned access may proceed, in which case raising the page-fault exception or access is more appropriate. Instruction address breakpoints have the same cause value as, but different priority than, data address breakpoints (a.k.a. watchpoints) and environment break exceptions (which are raised by the EBREAK instruction). Instruction address misaligned exceptions are raised by control-flow instructions with misaligned targets, rather than by the act of fetching an instruction. Therefore, these exceptions have lower priority than other instruction address exceptions.

3.1.16. Machine Trap Value Register (mtval)

The mtval register is an MXLEN-bit read-write register formatted as shown in Figure 24. When a trap is taken into M-mode, mtval is either set to zero or written with exception-specific information to assist software in handling the trap. Otherwise, mtval is never written by the implementation, though it may be explicitly written by software. The hardware platform will specify which exceptions must set mtval informatively and which may unconditionally set it to zero. If the hardware platform specifies that no exceptions set mtval to a nonzero value, then mtval is read-only zero.

If mtval is written with a nonzero value when a breakpoint, address-misaligned, access-fault, or page-fault exception occurs on an instruction fetch, load, or store, then mtval will contain the faulting virtual address.

When page-based virtual memory is enabled, mtval is written with the faulting virtual address, even for physical-memory access-fault exceptions. This design reduces datapath cost for most implementations, particularly those with hardware page-table walkers.

Figure 24. Machine Trap Value register.

If mtval is written with a nonzero value when a misaligned load or store causes an access-fault or page-fault exception, then mtval will contain the virtual address of the portion of the access that caused the fault.

If mtval is written with a nonzero value when an instruction access-fault or page-fault exception occurs on a system with variable-length instructions, then mtval will contain the virtual address of the portion of the instruction that caused the fault, while mepc will point to the beginning of the instruction.

The mtval register can optionally also be used to return the faulting instruction bits on an illegal-instruction exception (mepc points to the faulting instruction in memory). If mtval is written with a nonzero value when an illegal-instruction exception occurs, then mtval will contain the shortest of:

• the actual faulting instruction

• the first ILEN bits of the faulting instruction

• the first MXLEN bits of the faulting instruction

The value loaded into mtval on an illegal-instruction exception is right-justified and all unused upper bits are cleared to zero.

Capturing the faulting instruction in mtval reduces the overhead of instruction emulation, potentially avoiding several partial instruction loads if the instruction is misaligned, and likely data cache misses or slow uncached accesses when loads are used to fetch the instruction into a data register. There is also a problem of atomicity if another agent is manipulating the instruction memory, as might occur in a dynamic translation system. A requirement is that the entire instruction (or at least the first MXLEN bits) are fetched into mtval before taking the trap. This should not constrain implementations, which would typically fetch the entire instruction before attempting to decode the instruction, and avoids complicating software handlers. A value of zero in mtval signifies either that the feature is not supported, or an illegal zero instruction was fetched. A load from the instruction memory pointed to by mepc can be used to distinguish these two cases (or alternatively, the system configuration information can be interrogated to install the appropriate trap handling before runtime).

For other traps, mtval is set to zero, but a future standard may redefine mtval’s setting for other traps.

If mtval is not read-only zero, it is a WARL register that must be able to hold all valid virtual addresses and the value zero. It need not be capable of holding all possible invalid addresses. Prior to writing mtval, implementations may convert an invalid address into some other invalid address that mtval is capable of holding. If the feature to return the faulting instruction bits is implemented, mtval must also be able to hold all values less than 2^N, where N is the smaller of MXLEN and ILEN.

3.1.17. Machine Configuration Pointer Register (mconfigptr)

mconfigptr is an MXLEN-bit read-only CSR, formatted as shown in Figure 25, that holds the physical address of a configuration data structure. Software can traverse this data structure to discover information about the harts, the platform, and their configuration.

Figure 25. Machine Configuration Pointer register.

The pointer alignment in bits must be no smaller than the greatest supported MXLEN: i.e., if the greatest supported MXLEN is , then mconfigptr[-1:0] must be zero.

mconfigptr must be implemented, but it may be zero to indicate the configuration data structure does not exist or that an alternative mechanism must be used to locate it.

The format and schema of the configuration data structure have yet to be standardized. While mconfigptr will simply be hardwired in some implementations, other implementations may provide a means to configure the value returned on CSR reads. For example, mconfigptr might present the value of a memory-mapped register that is programmed by the platform or by M-mode software towards the beginning of the boot process.

3.1.18. Machine Environment Configuration Register (menvcfg)

The menvcfg CSR is a 64-bit read/write register, formatted as shown in , that controls certain characteristics of the execution environment for modes less privileged than M.

Figure 26. Machine environment configuration register (menvcfg).

If bit FIOM (Fence of I/O implies Memory) is set to one in menvcfg, FENCE instructions executed in modes less privileged than M are modified so the requirement to order accesses to device I/O implies also the requirement to order main memory accesses. Table 16 details the modified interpretation of FENCE instruction bits PI, PO, SI, and SO for modes less privileged than M when FIOM=1.

Similarly, for modes less privileged than M when FIOM=1, if an atomic instruction that accesses a region ordered as device I/O has its aq and/or rl bit set, then that instruction is ordered as though it accesses both device I/O and memory.

If S-mode is not supported, or if satp.MODE is read-only zero (always Bare), the implementation may make FIOM read-only zero.

Table 16. Modified interpretation of FENCE predecessor and successor sets for modes less privileged than M when FIOM=1.

Instruction bit	Meaning when set
PI PO	Predecessor device input and memory reads (PR implied) Predecessor device output and memory writes (PW implied)
SI SO	Successor device input and memory reads (SR implied) Successor device output and memory writes (SW implied)

Table 16. Modified interpretation of FENCE predecessor and successor sets for modes less privileged than M when FIOM=1. {#menvcfg-FIOM}

Bit FIOM is needed in menvcfg so M-mode can emulate the hypervisor extension of Chapter 9, which has an equivalent FIOM bit in the hypervisor CSR henvcfg.

The PBMTE bit controls whether the Svpbmt extension is available for use in S-mode and G-stage address translation (i.e., for page tables pointed to by satp or hgatp). When PBMTE=1, Svpbmt is available for S-mode and G-stage address translation. When PBMTE=0, the implementation behaves as though Svpbmt were not implemented. If Svpbmt is not implemented, PBMTE is read-only zero. Furthermore, for implementations with the hypervisor extension, henvcfg.PBMTE is read-only zero if menvcfg.PBMTE is zero.

The definition of the STCE field will be furnished by the forthcoming Sstc extension. Its allocation within menvcfg may change prior to the ratification of that extension.

The definition of the CBZE field will be furnished by the forthcoming Zicboz extension. Its allocation within menvcfg may change prior to the ratification of that extension.

The definitions of the CBCFE and CBIE fields will be furnished by the forthcoming Zicbom extension. Their allocations within menvcfg may change prior to the ratification of that extension.

When XLEN=32, menvcfgh is a 32-bit read/write register that aliases bits 63:32 of menvcfg. Register menvcfgh does not exist when XLEN=64.

If U-mode is not supported, then registers menvcfg and menvcfgh do not exist.

3.1.19. Machine Security Configuration Register (mseccfg)

mseccfg is an optional 64-bit read/write register, formatted as shown in Figure 27, that controls security features.

Figure 27. Machine security configuration register (mseccfg).

The definitions of the SSEED and USEED fields will be furnished by the forthcoming entropy-source extension, Zkr. Their allocations within mseccfg may change prior to the ratification of that extension.

The definitions of the RLB, MMWP, and MML fields will be furnished by the forthcoming PMP-enhancement extension, Smepmp. Their allocations within mseccfg may change prior to the ratification of that extension.

When XLEN=32 only, mseccfgh is a 32-bit read/write register that aliases bits 63:32 of mseccfg. Register mseccfgh does not exist when XLEN=64.

3.2. Machine-Level Memory-Mapped Registers

3.2.1. Machine Timer Registers (mtime and mtimecmp)

Platforms provide a real-time counter, exposed as a memory-mapped machine-mode read-write register, mtime. mtime must increment at constant frequency, and the platform must provide a mechanism for determining the period of an mtime tick. The mtime register will wrap around if the count overflows.

The mtime register has a 64-bit precision on all RV32 and RV64 systems. Platforms provide a 64-bit memory-mapped machine-mode timer compare register (mtimecmp). A machine timer interrupt becomes pending whenever mtime contains a value greater than or equal to mtimecmp, treating the values as unsigned integers. The interrupt remains posted until mtimecmp becomes greater than mtime (typically as a result of writing mtimecmp). The interrupt will only be taken if interrupts are enabled and the MTIE bit is set in the mie register.

Figure 28. Machine time register (memory-mapped control register).

Figure 29. Machine time compare register (memory-mapped control register).

The timer facility is defined to use wall-clock time rather than a cycle counter to support modern processors that run with a highly variable clock frequency to save energy through dynamic voltage and frequency scaling. Accurate real-time clocks (RTCs) are relatively expensive to provide (requiring a crystal or MEMS oscillator) and have to run even when the rest of system is powered down, and so there is usually only one in a system located in a different frequency/voltage domain from the processors. Hence, the RTC must be shared by all the harts in a system and accesses to the RTC will potentially incur the penalty of a voltage-level-shifter and clock-domain crossing. It is thus more natural to expose mtime as a memory-mapped register than as a CSR. Lower privilege levels do not have their own timecmp registers. Instead, machine-mode software can implement any number of virtual timers on a hart by multiplexing the next timer interrupt into the mtimecmp register. Simple fixed-frequency systems can use a single clock for both cycle counting and wall-clock time.

Writes to mtime and mtimecmp are guaranteed to be reflected in MTIP eventually, but not necessarily immediately.

A spurious timer interrupt might occur if an interrupt handler increments mtimecmp then immediately returns, because MTIP might not yet have fallen in the interim. All software should be written to assume this event is possible, but most software should assume this event is extremely unlikely. It is almost always more performant to incur an occasional spurious timer interrupt than to poll MTIP until it falls.

In RV32, memory-mapped writes to mtimecmp modify only one 32-bit part of the register. The following code sequence sets a 64-bit mtimecmp value without spuriously generating a timer interrupt due to the intermediate value of the comparand:

For RV64, naturally aligned 64-bit memory accesses to the mtime and mtimecmp registers are additionally supported and are atomic.

Sample code for setting the 64-bit time comparand in RV32 assuming a little-endian memory system and that the registers live in a strongly ordered I/O region. Storing -1 to the low-order bits of mtimecmp prevents mtimecmp from temporarily becoming smaller than the lesser of the old and new values.

            # New comparand is in a1:a0.
            li t0, -1
            la t1, mtimecmp
            sw t0, 0(t1)     # No smaller than old value.
            sw a1, 4(t1)     # No smaller than new value.
            sw a0, 0(t1)     # New value.

3.3. Machine-Mode Privileged Instructions

3.3.1. Environment Call and Breakpoint

The ECALL instruction is used to make a request to the supporting execution environment. When executed in U-mode, S-mode, or M-mode, it generates an environment-call-from-U-mode exception, environment-call-from-S-mode exception, or environment-call-from-M-mode exception, respectively, and performs no other operation.

ECALL generates a different exception for each originating privilege mode so that environment call exceptions can be selectively delegated. A typical use case for Unix-like operating systems is to delegate to S-mode the environment-call-from-U-mode exception but not the others.

The EBREAK instruction is used by debuggers to cause control to be transferred back to a debugging environment. It generates a breakpoint exception and performs no other operation.

As described in the "C" Standard Extension for Compressed Instructions in Volume I of this manual, the C.EBREAK instruction performs the same operation as the EBREAK instruction.

ECALL and EBREAK cause the receiving privilege mode’s epc register to be set to the address of the ECALL or EBREAK instruction itself, not the address of the following instruction. As ECALL and EBREAK cause synchronous exceptions, they are not considered to retire, and should not increment the minstret CSR.

3.3.2. Trap-Return Instructions

Instructions to return from trap are encoded under the PRIV minor opcode.

To return after handling a trap, there are separate trap return instructions per privilege level, MRET and SRET. MRET is always provided. SRET must be provided if supervisor mode is supported, and should raise an illegal-instruction exception otherwise. SRET should also raise an illegal-instruction exception when TSR=1 in mstatus, as described in Section 3.1.6.5. An xRET instruction can be executed in privilege mode x or higher, where executing a lower-privilege xRET instruction will pop the relevant lower-privilege interrupt enable and privilege mode stack. In addition to manipulating the privilege stack as described in Section 3.1.6.1, xRET sets the pc to the value stored in the xepc register.

If the A extension is supported, the xRET instruction is allowed to clear any outstanding LR address reservation but is not required to. Trap handlers should explicitly clear the reservation if required (e.g., by using a dummy SC) before executing the xRET.

If xRET instructions always cleared LR reservations, it would be impossible to single-step through LR/SC sequences using a debugger.

3.3.3. Wait for Interrupt

The Wait for Interrupt instruction (WFI) informs the implementation that the current hart can be stalled until an interrupt might need servicing. Execution of the WFI instruction can also be used to inform the hardware platform that suitable interrupts should preferentially be routed to this hart. WFI is available in all privileged modes, and optionally available to U-mode. This instruction may raise an illegal-instruction exception when TW=1 in mstatus, as described in Section 3.1.6.5.

If an enabled interrupt is present or later becomes present while the hart is stalled, the interrupt trap will be taken on the following instruction, i.e., execution resumes in the trap handler and mepc = pc + 4.

The following instruction takes the interrupt trap so that a simple return from the trap handler will execute code after the WFI instruction.

Implementations are permitted to resume execution for any reason, even if an enabled interrupt has not become pending. Hence, a legal implementation is to simply implement the WFI instruction as a NOP.

If the implementation does not stall the hart on execution of the instruction, then the interrupt will be taken on some instruction in the idle loop containing the WFI, and on a simple return from the handler, the idle loop will resume execution.

The WFI instruction can also be executed when interrupts are disabled. The operation of WFI must be unaffected by the global interrupt bits in mstatus (MIE and SIE) and the delegation register mideleg (i.e., the hart must resume if a locally enabled interrupt becomes pending, even if it has been delegated to a less-privileged mode), but should honor the individual interrupt enables (e.g, MTIE) (i.e., implementations should avoid resuming the hart if the interrupt is pending but not individually enabled). WFI is also required to resume execution for locally enabled interrupts pending at any privilege level, regardless of the global interrupt enable at each privilege level.

If the event that causes the hart to resume execution does not cause an interrupt to be taken, execution will resume at pc + 4, and software must determine what action to take, including looping back to repeat the WFI if there was no actionable event.

By allowing wakeup when interrupts are disabled, an alternate entry point to an interrupt handler can be called that does not require saving the current context, as the current context can be saved or discarded before the WFI is executed. As implementations are free to implement WFI as a NOP, software must explicitly check for any relevant pending but disabled interrupts in the code following an WFI, and should loop back to the WFI if no suitable interrupt was detected. The mip or sip registers can be interrogated to determine the presence of any interrupt in machine or supervisor mode respectively. The operation of WFI is unaffected by the delegation register settings. WFI is defined so that an implementation can trap into a higher privilege mode, either immediately on encountering the WFI or after some interval to initiate a machine-mode transition to a lower power state, for example. The same "wait-for-event" template might be used for possible future extensions that wait on memory locations changing, or message arrival.

3.3.4. Custom SYSTEM Instructions

The subspace of the SYSTEM major opcode shown in Figure 30 is designated for custom use. It is recommended that these instructions use bits 29:28 to designate the minimum required privilege mode, as do other SYSTEM instructions.

Figure 30. SYSTEM instruction encodings designated for custom use.

3.4. Reset

Upon reset, a hart’s privilege mode is set to M. The mstatus fields MIE and MPRV are reset to 0. If little-endian memory accesses are supported, the mstatus/mstatush field MBE is reset to 0. The misa register is reset to enable the maximal set of supported extensions and widest MXLEN, as described in Section 3.1.1. For implementations with the "A" standard extension, there is no valid load reservation. The pc is set to an implementation-defined reset vector. The mcause register is set to a value indicating the cause of the reset. Writable PMP registers’ A and L fields are set to 0, unless the platform mandates a different reset value for some PMP registers’ A and L fields. If the hypervisor extension is implemented, the hgatp.MODE and vsatp.MODE fields are reset to 0. If the Smrnmi extension is implemented, the mnstatus.NMIE field is reset to 0. No WARL field contains an illegal value. All other hart state is UNSPECIFIED.

The mcause values after reset have implementation-specific interpretation, but the value 0 should be returned on implementations that do not distinguish different reset conditions. Implementations that distinguish different reset conditions should only use 0 to indicate the most complete reset.

Some designs may have multiple causes of reset (e.g., power-on reset, external hard reset, brownout detected, watchdog timer elapse, sleep-mode wakeup), which machine-mode software and debuggers may wish to distinguish. mcause reset values may alias mcause values following synchronous exceptions. There should be no ambiguity in this overlap, since on reset the pc is typically set to a different value than on other traps.

3.5. Non-Maskable Interrupts

Non-maskable interrupts (NMIs) are only used for hardware error conditions, and cause an immediate jump to an implementation-defined NMI vector running in M-mode regardless of the state of a hart’s interrupt enable bits. The mepc register is written with the virtual address of the instruction that was interrupted, and mcause is set to a value indicating the source of the NMI. The NMI can thus overwrite state in an active machine-mode interrupt handler.

The values written to mcause on an NMI are implementation-defined. The high Interrupt bit of mcause should be set to indicate that this was an interrupt. An Exception Code of 0 is reserved to mean "unknown cause" and implementations that do not distinguish sources of NMIs via the mcause register should return 0 in the Exception Code.

Unlike resets, NMIs do not reset processor state, enabling diagnosis, reporting, and possible containment of the hardware error.

3.6. Physical Memory Attributes

The physical memory map for a complete system includes various address ranges, some corresponding to memory regions, some to memory-mapped control registers, and some to vacant holes in the address space. Some memory regions might not support reads, writes, or execution; some might not support subword or subblock accesses; some might not support atomic operations; and some might not support cache coherence or might have different memory models. Similarly, memory-mapped control registers vary in their supported access widths, support for atomic operations, and whether read and write accesses have associated side effects. In RISC-V systems, these properties and capabilities of each region of the machine’s physical address space are termed physical memory attributes (PMAs). This section describes RISC-V PMA terminology and how RISC-V systems implement and check PMAs.

PMAs are inherent properties of the underlying hardware and rarely change during system operation. Unlike physical memory protection values described in Section 3.7, PMAs do not vary by execution context. The PMAs of some memory regions are fixed at chip design time—for example, for an on-chip ROM. Others are fixed at board design time, depending, for example, on which other chips are connected to off-chip buses. Off-chip buses might also support devices that could be changed on every power cycle (cold pluggable) or dynamically while the system is running (hot pluggable). Some devices might be configurable at run time to support different uses that imply different PMAs—for example, an on-chip scratchpad RAM might be cached privately by one core in one end-application, or accessed as a shared non-cached memory in another end-application.

Most systems will require that at least some PMAs are dynamically checked in hardware later in the execution pipeline after the physical address is known, as some operations will not be supported at all physical memory addresses, and some operations require knowing the current setting of a configurable PMA attribute. While many other architectures specify some PMAs in the virtual memory page tables and use the TLB to inform the pipeline of these properties, this approach injects platform-specific information into a virtualized layer and can cause system errors unless attributes are correctly initialized in each page-table entry for each physical memory region. In addition, the available page sizes might not be optimal for specifying attributes in the physical memory space, leading to address-space fragmentation and inefficient use of expensive TLB entries.

For RISC-V, we separate out specification and checking of PMAs into a separate hardware structure, the PMA checker. In many cases, the attributes are known at system design time for each physical address region, and can be hardwired into the PMA checker. Where the attributes are run-time configurable, platform-specific memory-mapped control registers can be provided to specify these attributes at a granularity appropriate to each region on the platform (e.g., for an on-chip SRAM that can be flexibly divided between cacheable and uncacheable uses). PMAs are checked for any access to physical memory, including accesses that have undergone virtual to physical memory translation. To aid in system debugging, we strongly recommend that, where possible, RISC-V processors precisely trap physical memory accesses that fail PMA checks. Precisely trapped PMA violations manifest as instruction, load, or store access-fault exceptions, distinct from virtual-memory page-fault exceptions. Precise PMA traps might not always be possible, for example, when probing a legacy bus architecture that uses access failures as part of the discovery mechanism. In this case, error responses from peripheral devices will be reported as imprecise bus-error interrupts.

PMAs must also be readable by software to correctly access certain devices or to correctly configure other hardware components that access memory, such as DMA engines. As PMAs are tightly tied to a given physical platform’s organization, many details are inherently platform-specific, as is the means by which software can learn the PMA values for a platform. Some devices, particularly legacy buses, do not support discovery of PMAs and so will give error responses or time out if an unsupported access is attempted. Typically, platform-specific machine-mode code will extract PMAs and ultimately present this information to higher-level less-privileged software using some standard representation.

Where platforms support dynamic reconfiguration of PMAs, an interface will be provided to set the attributes by passing requests to a machine-mode driver that can correctly reconfigure the platform. For example, switching cacheability attributes on some memory regions might involve platform-specific operations, such as cache flushes, that are available only to machine-mode.

3.6.1. Main Memory versus I/O versus Vacant Regions

The most important characterization of a given memory address range is whether it holds regular main memory, or I/O devices, or is vacant. Regular main memory is required to have a number of properties, specified below, whereas I/O devices can have a much broader range of attributes. Memory regions that do not fit into regular main memory, for example, device scratchpad RAMs, are categorized as I/O regions. Vacant regions are also classified as I/O regions but with attributes specifying that no accesses are supported.

3.6.2. Supported Access Type PMAs

Access types specify which access widths, from 8-bit byte to long multi-word burst, are supported, and also whether misaligned accesses are supported for each access width.

Although software running on a RISC-V hart cannot directly generate bursts to memory, software might have to program DMA engines to access I/O devices and might therefore need to know which access sizes are supported.

Main memory regions always support read and write of all access widths required by the attached devices, and can specify whether instruction fetch is supported.

Some platforms might mandate that all of main memory support instruction fetch. Other platforms might prohibit instruction fetch from some main memory regions. In some cases, the design of a processor or device accessing main memory might support other widths, but must be able to function with the types supported by the main memory.

I/O regions can specify which combinations of read, write, or execute accesses to which data widths are supported.

For systems with page-based virtual memory, I/O and memory regions can specify which combinations of hardware page-table reads and hardware page-table writes are supported.

Unix-like operating systems generally require that all of cacheable main memory supports page-table walks.

3.6.3. Atomicity PMAs

Atomicity PMAs describes which atomic instructions are supported in this address region. Support for atomic instructions is divided into two categories: LR/SC and AMOs.

Some platforms might mandate that all of cacheable main memory support all atomic operations required by the attached processors.

3.6.3.1. AMO PMA

Within AMOs, there are four levels of support: AMONone, AMOSwap, AMOLogical, and AMOArithmetic. AMONone indicates that no AMO operations are supported. AMOSwap indicates that only amoswap instructions are supported in this address range. AMOLogical indicates that swap instructions plus all the logical AMOs (amoand, amoor, amoxor) are supported. AMOArithmetic indicates that all RISC-V AMOs are supported. For each level of support, naturally aligned AMOs of a given width are supported if the underlying memory region supports reads and writes of that width. Main memory and I/O regions may only support a subset or none of the processor-supported atomic operations.

Table 17. Classes of AMOs supported by I/O regions.

AMO Class	Supported Operations
AMONone AMOSwap AMOLogical AMOArithmetic	None amoswap above + amoand, amoor, amoxor above + amoadd, amomin, amomax, amominu, amomaxu

Table 17. Classes of AMOs supported by I/O regions.

We recommend providing at least AMOLogical support for I/O regions where possible.

3.6.3.2. Reservability PMA

For LR/SC, there are three levels of support indicating combinations of the reservability and eventuality properties: RsrvNone, RsrvNonEventual, and RsrvEventual. RsrvNone indicates that no LR/SC operations are supported (the location is non-reservable). RsrvNonEventual indicates that the operations are supported (the location is reservable), but without the eventual success guarantee described in the unprivileged ISA specification. RsrvEventual indicates that the operations are supported and provide the eventual success guarantee.

We recommend providing RsrvEventual support for main memory regions where possible. Most I/O regions will not support LR/SC accesses, as these are most conveniently built on top of a cache-coherence scheme, but some may support RsrvNonEventual or RsrvEventual. When LR/SC is used for memory locations marked RsrvNonEventual, software should provide alternative fall-back mechanisms used when lack of progress is detected.

3.6.3.3. Alignment

Memory regions that support aligned LR/SC or aligned AMOs might also support misaligned LR/SC or misaligned AMOs for some addresses and access widths. If, for a given address and access width, a misaligned LR/SC or AMO generates an address-misaligned exception, then all loads, stores, LRs/SCs, and AMOs using that address and access width must generate address-misaligned exceptions.

The standard "A" extension does not support misaligned AMOs or LR/SC pairs. Support for misaligned AMOs is provided by the standard "Zam" extension. Support for misaligned LR/SC sequences is not currently standardized, so LR and SC to misaligned addresses must raise an exception. Mandating that misaligned loads and stores raise address-misaligned exceptions wherever misaligned AMOs raise address-misaligned exceptions permits the emulation of misaligned AMOs in an M-mode trap handler. The handler guarantees atomicity by acquiring a global mutex and emulating the access within the critical section. Provided that the handler for misaligned loads and stores uses the same mutex, all accesses to a given address that use the same word size will be mutually atomic.

Implementations may raise access-fault exceptions instead of address-misaligned exceptions for some misaligned accesses, indicating the instruction should not be emulated by a trap handler. If, for a given address and access width, all misaligned LRs/SCs and AMOs generate access-fault exceptions, then regular misaligned loads and stores using the same address and access width are not required to execute atomically.

3.6.4. Memory-Ordering PMAs

Regions of the address space are classified as either main memory or I/O for the purposes of ordering by the FENCE instruction and atomic-instruction ordering bits.

Accesses by one hart to main memory regions are observable not only by other harts but also by other devices with the capability to initiate requests in the main memory system (e.g., DMA engines). Coherent main memory regions always have either the RVWMO or RVTSO memory model. Incoherent main memory regions have an implementation-defined memory model.

Accesses by one hart to an I/O region are observable not only by other harts and bus mastering devices but also by the targeted I/O devices, and I/O regions may be accessed with either relaxed or strong ordering. Accesses to an I/O region with relaxed ordering are generally observed by other harts and bus mastering devices in a manner similar to the ordering of accesses to an RVWMO memory region, as discussed in Section A.4.2 in Volume I of this specification. By contrast, accesses to an I/O region with strong ordering are generally observed by other harts and bus mastering devices in program order.

Each strongly ordered I/O region specifies a numbered ordering channel, which is a mechanism by which ordering guarantees can be provided between different I/O regions. Channel 0 is used to indicate point-to-point strong ordering only, where only accesses by the hart to the single associated I/O region are strongly ordered.

Channel 1 is used to provide global strong ordering across all I/O regions. Any accesses by a hart to any I/O region associated with channel 1 can only be observed to have occurred in program order by all other harts and I/O devices, including relative to accesses made by that hart to relaxed I/O regions or strongly ordered I/O regions with different channel numbers. In other words, any access to a region in channel 1 is equivalent to executing a fence io,io instruction before and after the instruction.

Other larger channel numbers provide program ordering to accesses by that hart across any regions with the same channel number.

Systems might support dynamic configuration of ordering properties on each memory region.

Strong ordering can be used to improve compatibility with legacy device driver code, or to enable increased performance compared to insertion of explicit ordering instructions when the implementation is known to not reorder accesses. Local strong ordering (channel 0) is the default form of strong ordering as it is often straightforward to provide if there is only a single in-order communication path between the hart and the I/O device. Generally, different strongly ordered I/O regions can share the same ordering channel without additional ordering hardware if they share the same interconnect path and the path does not reorder requests.

3.6.5. Coherence and Cacheability PMAs

Coherence is a property defined for a single physical address, and indicates that writes to that address by one agent will eventually be made visible to other agents in the system. Coherence is not to be confused with the memory consistency model of a system, which defines what values a memory read can return given the previous history of reads and writes to the entire memory system. In RISC-V platforms, the use of hardware-incoherent regions is discouraged due to software complexity, performance, and energy impacts.

The cacheability of a memory region should not affect the software view of the region except for differences reflected in other PMAs, such as main memory versus I/O classification, memory ordering, supported accesses and atomic operations, and coherence. For this reason, we treat cacheability as a platform-level setting managed by machine-mode software only.

Where a platform supports configurable cacheability settings for a memory region, a platform-specific machine-mode routine will change the settings and flush caches if necessary, so the system is only incoherent during the transition between cacheability settings. This transitory state should not be visible to lower privilege levels.

Coherence is straightforward to provide for a shared memory region that is not cached by any agent. The PMA for such a region would simply indicate it should not be cached in a private or shared cache. Coherence is also straightforward for read-only regions, which can be safely cached by multiple agents without requiring a cache-coherence scheme. The PMA for this region would indicate that it can be cached, but that writes are not supported. Some read-write regions might only be accessed by a single agent, in which case they can be cached privately by that agent without requiring a coherence scheme. The PMA for such regions would indicate they can be cached. The data can also be cached in a shared cache, as other agents should not access the region. If an agent can cache a read-write region that is accessible by other agents, whether caching or non-caching, a cache-coherence scheme is required to avoid use of stale values. In regions lacking hardware cache coherence (hardware-incoherent regions), cache coherence can be implemented entirely in software, but software coherence schemes are notoriously difficult to implement correctly and often have severe performance impacts due to the need for conservative software-directed cache-flushing. Hardware cache-coherence schemes require more complex hardware and can impact performance due to the cache-coherence probes, but are otherwise invisible to software. For each hardware cache-coherent region, the PMA would indicate that the region is coherent and which hardware coherence controller to use if the system has multiple coherence controllers. For some systems, the coherence controller might be an outer-level shared cache, which might itself access further outer-level cache-coherence controllers hierarchically. Most memory regions within a platform will be coherent to software, because they will be fixed as either uncached, read-only, hardware cache-coherent, or only accessed by one agent.

If a PMA indicates non-cacheability, then accesses to that region must be satisfied by the memory itself, not by any caches.

For implementations with a cacheability-control mechanism, the situation may arise that a program uncacheably accesses a memory location that is currently cache-resident. In this situation, the cached copy must be ignored. This constraint is necessary to prevent more-privileged modes’ speculative cache refills from affecting the behavior of less-privileged modes’ uncacheable accesses.

3.6.6. Idempotency PMAs

Idempotency PMAs describe whether reads and writes to an address region are idempotent. Main memory regions are assumed to be idempotent. For I/O regions, idempotency on reads and writes can be specified separately (e.g., reads are idempotent but writes are not). If accesses are non-idempotent, i.e., there is potentially a side effect on any read or write access, then speculative or redundant accesses must be avoided.

For the purposes of defining the idempotency PMAs, changes in observed memory ordering created by redundant accesses are not considered a side effect.

While hardware should always be designed to avoid speculative or redundant accesses to memory regions marked as non-idempotent, it is also necessary to ensure software or compiler optimizations do not generate spurious accesses to non-idempotent memory regions. Non-idempotent regions might not support misaligned accesses. Misaligned accesses to such regions should raise access-fault exceptions rather than address-misaligned exceptions, indicating that software should not emulate the misaligned access using multiple smaller accesses, which could cause unexpected side effects.

For non-idempotent regions, implicit reads and writes must not be performed early or speculatively, with the following exceptions. When a non-speculative implicit read is performed, an implementation is permitted to additionally read any of the bytes within a naturally aligned power-of-2 region containing the address of the non-speculative implicit read. Furthermore, when a non-speculative instruction fetch is performed, an implementation is permitted to additionally read any of the bytes within the next naturally aligned power-of-2 region of the same size (with the address of the region taken modulo 2^XLEN. The results of these additional reads may be used to satisfy subsequent early or speculative implicit reads. The size of these naturally aligned power-of-2 regions is implementation-defined, but, for systems with page-based virtual memory, must not exceed the smallest supported page size.

3.7. Physical Memory Protection

To support secure processing and contain faults, it is desirable to limit the physical addresses accessible by software running on a hart. An optional physical memory protection (PMP) unit provides per-hart machine-mode control registers to allow physical memory access privileges (read, write, execute) to be specified for each physical memory region. The PMP values are checked in parallel with the PMA checks described in Section 3.6.

The granularity of PMP access control settings are platform-specific, but the standard PMP encoding supports regions as small as four bytes. Certain regions’ privileges can be hardwired—for example, some regions might only ever be visible in machine mode but in no lower-privilege layers.

Platforms vary widely in demands for physical memory protection, and some platforms may provide other PMP structures in addition to or instead of the scheme described in this section.

PMP checks are applied to all accesses whose effective privilege mode is S or U, including instruction fetches and data accesses in S and U mode, and data accesses in M-mode when the MPRV bit in mstatus is set and the MPP field in mstatus contains S or U. PMP checks are also applied to page-table accesses for virtual-address translation, for which the effective privilege mode is S. Optionally, PMP checks may additionally apply to M-mode accesses, in which case the PMP registers themselves are locked, so that even M-mode software cannot change them until the hart is reset. In effect, PMP can grant permissions to S and U modes, which by default have none, and can revoke permissions from M-mode, which by default has full permissions.

PMP violations are always trapped precisely at the processor.

3.7.1. Physical Memory Protection CSRs

PMP entries are described by an 8-bit configuration register and one MXLEN-bit address register. Some PMP settings additionally use the address register associated with the preceding PMP entry. Up to 64 PMP entries are supported. Implementations may implement zero, 16, or 64 PMP entries; the lowest-numbered PMP entries must be implemented first. All PMP CSR fields are WARL and may be read-only zero. PMP CSRs are only accessible to M-mode.

The PMP configuration registers are densely packed into CSRs to minimize context-switch time. For RV32, sixteen CSRs, pmpcfg0–pmpcfg15, hold the configurations pmp0cfg–pmp63cfg for the 64 PMP entries, as shown in Figure 31. For RV64, eight even-numbered CSRs, pmpcfg0, pmpcfg2, …, pmpcfg14, hold the configurations for the 64 PMP entries, as shown in Figure 32. For RV64, the odd-numbered configuration registers, pmpcfg1, pmpcfg3, …, pmpcfg15, are illegal.

RV64 systems use pmpcfg2, rather than pmpcfg1, to hold configurations for PMP entries 8-15. This design reduces the cost of supporting multiple MXLEN values, since the configurations for PMP entries 8-11 appear in pmpcfg2[31:0] for both RV32 and RV64.

Figure 31. RV32 PMP configuration CSR layout.

Figure 32. RV64 PMP configuration CSR layout.

The PMP address registers are CSRs named pmpaddr0-pmpaddr63. Each PMP address register encodes bits 33-2 of a 34-bit physical address for RV32, as shown in Figure 33. For RV64, each PMP address register encodes bits 55-2 of a 56-bit physical address, as shown in Figure 34. Not all physical address bits may be implemented, and so the pmpaddr registers are WARL.

The Sv32 page-based virtual-memory scheme described in Section 5.3 supports 34-bit physical addresses for RV32, so the PMP scheme must support addresses wider than XLEN for RV32. The Sv39 and Sv48 page-based virtual-memory schemes described in Section 5.4 and Section 5.5 support a 56-bit physical address space, so the RV64 PMP address registers impose the same limit.

Figure 33. PMP address register format, RV32.

Figure 34. PMP address register format, RV64.

Figure 35 shows the layout of a PMP configuration register. The R, W, and X bits, when set, indicate that the PMP entry permits read, write, and instruction execution, respectively. When one of these bits is clear, the corresponding access type is denied. The R, W, and X fields form a collective WARL field for which the combinations with R=0 and W=1 are reserved. The remaining two fields, A and L, are described in the following sections.

Figure 35. PMP configuration register format.

Attempting to fetch an instruction from a PMP region that does not have execute permissions raises an instruction access-fault exception. Attempting to execute a load or load-reserved instruction which accesses a physical address within a PMP region without read permissions raises a load access-fault exception. Attempting to execute a store, store-conditional, or AMO instruction which accesses a physical address within a PMP region without write permissions raises a store access-fault exception.

If MXLEN is changed, the contents of the pmpxcfg fields are preserved, but appear in the pmpcfgy CSR prescribed by the new setting of MXLEN. For example, when MXLEN is changed from 64 to 32, pmp4cfg moves from pmpcfg0[39:32] to pmpcfg1[7:0]. The pmpaddr CSRs follow the usual CSR width modulation rules described in Section 2.6.

3.7.1.1. Address Matching

The A field in a PMP entry’s configuration register encodes the address-matching mode of the associated PMP address register. The encoding of this field is shown in Table 18. When A=0, this PMP entry is disabled and matches no addresses. Two other address-matching modes are supported: naturally aligned power-of-2 regions (NAPOT), including the special case of naturally aligned four-byte regions (NA4); and the top boundary of an arbitrary range (TOR). These modes support four-byte granularity.

Table 18. Encoding of A field in PMP configuration registers.

A	Name	Description
0 1 2 3	OFF TOR NA4 NAPOT	Null region (disabled) Top of range Naturally aligned four-byte region Naturally aligned power-of-two region, ≥8 bytes

Table 18. Encoding of A field in PMP configuration registers. {#pmpcfg-a}

NAPOT ranges make use of the low-order bits of the associated address register to encode the size of the range, as shown in Table 19.

Table 19. NAPOT range encoding in PMP address and configuration registers.

pmpaddr	pmpcfg.A	Match type and size
yyyy…​yyyy yyyy…​yyy0 yyyy…​yy01 yyyy…​y011 …​ yy01…​1111 y011…​1111 0111…​1111 1111…​1111	NA4 NAPOT NAPOT NAPOT …​ NAPOT NAPOT NAPOT NAPOT	4-byte NAPOT range 8-byte NAPOT range 16-byte NAPOT range 32-byte NAPOT range … 2XLEN-byte NAPOT range 2XLEN+1-byte NAPOT range 2XLEN+2-byte NAPOT range 2XLEN+3-byte NAPOT range

Table 19. NAPOT range encoding in PMP address and configuration registers. {#pmpcfg-napot}

If TOR is selected, the associated address register forms the top of the address range, and the preceding PMP address register forms the bottom of the address range. If PMP entry i's A field is set to TOR, the entry matches any address y such that pmpaddri-1≤y<pmpaddri (irrespective of the value of pmpcfgi-1). If PMP entry 0’s A field is set to TOR, zero is used for the lower bound, and so it matches any address y<pmpaddr0.

If pmpaddri-1≥pmpaddri and pmpcfgi.A=TOR, then PMP entry i matches no addresses.

Although the PMP mechanism supports regions as small as four bytes, platforms may specify coarser PMP regions. In general, the PMP grain is bytes and must be the same across all PMP regions. When , the NA4 mode is not selectable. When and .A[1] is set, i.e. the mode is NAPOT, then bits [G-2:0] read as all ones. When and .A[1] is clear, i.e. the mode is OFF or TOR, then bits [G-1:0] read as all zeros. Bits [G-1:0] do not affect the TOR address-matching logic. Although changing .A[1] affects the value read from , it does not affect the underlying value stored in that register—in particular, [G-1] retains its original value when .A is changed from NAPOT to TOR/OFF then back to NAPOT.

Software may determine the PMP granularity by writing zero to pmp0cfg, then writing all ones to pmpaddr0, then reading back pmpaddr0. If G is the index of the least-significant bit set, the PMP granularity is 2G+2 bytes.

If the current XLEN is greater than MXLEN, the PMP address registers are zero-extended from MXLEN to XLEN bits for the purposes of address matching.

3.7.1.2. Locking and Privilege Mode

The L bit indicates that the PMP entry is locked, i.e., writes to the configuration register and associated address registers are ignored. Locked PMP entries remain locked until the hart is reset. If PMP entry i is locked, writes to pmpicfg and pmpaddri are ignored. Additionally, if PMP entry i is locked and pmpicfg.A is set to TOR, writes to pmpaddri-1 are ignored.

Setting the L bit locks the PMP entry even when the A field is set to OFF.

In addition to locking the PMP entry, the L bit indicates whether the R/W/X permissions are enforced on M-mode accesses. When the L bit is set, these permissions are enforced for all privilege modes. When the L bit is clear, any M-mode access matching the PMP entry will succeed; the R/W/X permissions apply only to S and U modes.

3.7.1.3. Priority and Matching Logic

PMP entries are statically prioritized. The lowest-numbered PMP entry that matches any byte of an access determines whether that access succeeds or fails. The matching PMP entry must match all bytes of an access, or the access fails, irrespective of the L, R, W, and X bits. For example, if a PMP entry is configured to match the four-byte range 0xC–0xF, then an 8-byte access to the range 0x8–0xF will fail, assuming that PMP entry is the highest-priority entry that matches those addresses.

If a PMP entry matches all bytes of an access, then the L, R, W, and X bits determine whether the access succeeds or fails. If the L bit is clear and the privilege mode of the access is M, the access succeeds. Otherwise, if the L bit is set or the privilege mode of the access is S or U, then the access succeeds only if the R, W, or X bit corresponding to the access type is set.

If no PMP entry matches an M-mode access, the access succeeds. If no PMP entry matches an S-mode or U-mode access, but at least one PMP entry is implemented, the access fails.

If at least one PMP entry is implemented, but all PMP entries’ A fields are set to OFF, then all S-mode and U-mode memory accesses will fail.

Failed accesses generate an instruction, load, or store access-fault exception. Note that a single instruction may generate multiple accesses, which may not be mutually atomic. An access-fault exception is generated if at least one access generated by an instruction fails, though other accesses generated by that instruction may succeed with visible side effects. Notably, instructions that reference virtual memory are decomposed into multiple accesses.

On some implementations, misaligned loads, stores, and instruction fetches may also be decomposed into multiple accesses, some of which may succeed before an access-fault exception occurs. In particular, a portion of a misaligned store that passes the PMP check may become visible, even if another portion fails the PMP check. The same behavior may manifest for stores wider than XLEN bits (e.g., the FSD instruction in RV32D), even when the store address is naturally aligned.

3.7.2. Physical Memory Protection and Paging

The Physical Memory Protection mechanism is designed to compose with the page-based virtual memory systems described in Chapter 5. When paging is enabled, instructions that access virtual memory may result in multiple physical-memory accesses, including implicit references to the page tables. The PMP checks apply to all of these accesses. The effective privilege mode for implicit page-table accesses is S.

Implementations with virtual memory are permitted to perform address translations speculatively and earlier than required by an explicit memory access, and are permitted to cache them in address translation cache structures—including possibly caching the identity mappings from effective address to physical address used in Bare translation modes and M-mode. The PMP settings for the resulting physical address may be checked (and possibly cached) at any point between the address translation and the explicit memory access. Hence, when the PMP settings are modified, M-mode software must synchronize the PMP settings with the virtual memory system and any PMP or address-translation caches. This is accomplished by executing an SFENCE.VMA instruction with rs1=x0 and rs2=x0, after the PMP CSRs are written.

If page-based virtual memory is not implemented, memory accesses check the PMP settings synchronously, so no SFENCE.VMA is needed.

4. "Smrnmi" Standard Extension for Resumable Non-Maskable Interrupts, Version 0.4

Warning! This draft specification may change before being accepted as standard by RISC-V International.

The base machine-level architecture supports only unresumable non-maskable interrupts (UNMIs), where the NMI jumps to a handler in machine mode, overwriting the current mepc and mcause register values. If the hart had been executing machine-mode code in a trap handler, the previous values in mepc and mcause would not be recoverable and so execution is not generally resumable.

The Smrnmi extension adds support for resumable non-maskable interrupts (RNMIs) to RISC-V. The extension adds four new CSRs (mnepc, mncause, mnstatus, and mnscratch) to hold the interrupted state, and one new instruction, MNRET, to resume from the RNMI handler.

4.1. RNMI Interrupt Signals

The rnmi interrupt signals are inputs to the hart. These interrupts have higher priority than any other interrupt or exception on the hart and cannot be disabled by software. Specifically, they are not disabled by clearing the mstatus.MIE register.

4.2. RNMI Handler Addresses

The RNMI interrupt trap handler address is implementation-defined.

RNMI also has an associated exception trap handler address, which is implementation defined.

4.3. RNMI CSRs

This proposal adds additional M-mode CSRs to enable a resumable non-maskable interrupt (RNMI).

Figure 36. Resumable NMI scratch register mnscratch

The mnscratch CSR holds an MXLEN-bit read-write register which enables the NMI trap handler to save and restore the context that was interrupted.

Figure 37. Resumable NMI program counter mnepc.

The mnepc CSR is an MXLEN-bit read-write register which on entry to the NMI trap handler holds the PC of the instruction that took the interrupt.

The low bit of mnepc (mnepc[0]) is always zero. On implementations that support only IALIGN=32, the two low bits (mnepc[1:0]) are always zero.

If an implementation allows IALIGN to be either 16 or 32 (by changing CSR misa, for example), then, whenever IALIGN=32, bit mnepc[1] is masked on reads so that it appears to be 0. This masking occurs also for the implicit read by the MRET instruction. Though masked, mnepc[1] remains writable when IALIGN=32.

mnepc is a WARL register that must be able to hold all valid virtual addresses. It need not be capable of holding all possible invalid addresses. Prior to writing mnepc, implementations may convert an invalid address into some other invalid address that mnepc is capable of holding.

Figure 38. Resumable NMI cause mncause.

The mncause CSR holds the reason for the NMI, with bit MXLEN-1 set to 1, and the NMI cause encoded in the least-significant bits or zero if NMI causes are not supported.

Figure 39. Resumable NMI status register mnstatus.

The mnstatus CSR holds a two-bit field, MNPP, which on entry to the trap handler holds the privilege mode of the interrupted context, encoded in the same manner as mstatus.MPP. It also holds a one-bit field, MNPV, which on entry to the trap handler holds the virtualization mode of the interrupted context, encoded in the same manner as mstatus.MPV.

mnstatus also holds the NMIE bit. When NMIE=1, nonmaskable interrupts are enabled. When NMIE=0, all interrupts are disabled.

When NMIE=0, the hart behaves as though mstatus.MPRV were clear, regardless of the current setting of mstatus.MPRV.

Upon reset, NMIE contains the value 0.

RNMIs are masked out of reset to give software the opportunity to initialize data structures and devices for subsequent RNMI handling.

Software can set NMIE to 1, but attempts to clear NMIE have no effect.

Normally, only reset sequences will explicitly set the NMIE bit. That the NMIE bit is settable does not suffice to support the nesting of RNMIs. To support this feature in a direct manner would have required allowing software to clear the NMIE bit—a design choice that would have contravened the concept of non-maskability. Software that wishes to minimize the latency until the next RNMI is taken can follow the top-half/bottom-half model, where the RNMI handler itself only enqueues a task to a task queue then returns. The bulk of the interrupt servicing is performed later, with RNMIs enabled.

For the purposes of the WFI instruction, NMIE is a global interrupt enable, meaning that the setting of NMIE does not affect the operation of the WFI instruction.

The other bits in mnstatus are reserved; software should write zeros and hardware implementations should return zeros.

4.4. MNRET Instruction

MNRET is an M-mode-only instruction that uses the values in mnepc and mnstatus to return to the program counter, privilege mode, and virtualization mode of the interrupted context. This instruction also sets mnstatus.NMIE. If MNRET changes the privilege mode to a mode less privileged than M, it also sets mstatus.MPRV to 0.

4.5. RNMI Operation

When an RNMI interrupt is detected, the interrupted PC is written to the mnepc CSR, the type of RNMI to the mncause CSR, and the privilege mode of the interrupted context to the mnstatus CSR. The mnstatus.NMIE bit is cleared, masking all interrupts.

The hart then enters machine-mode and jumps to the RNMI trap handler address.

The RNMI handler can resume original execution using the new MNRET instruction, which restores the PC from mnepc, the privilege mode from mnstatus, and also sets mnstatus.NMIE, which re-enables interrupts.

If the hart encounters an exception while executing in M-mode with the mnstatus.NMIE bit clear, the actions taken are the same as if the exception had occurred while mnstatus.NMIE were set, except that the program counter is set to the RNMI exception trap handler address (rather than the address specified by mtvec).

The Smrnmi extension does not change the behavior of the MRET and SRET instructions. In particular, MRET and SRET are unaffected by the mnstatus.NMIE bit, and their execution does not alter the mnstatus.NMIE bit.

5. Supervisor-Level ISA, Version 1.12

This chapter describes the RISC-V supervisor-level architecture, which contains a common core that is used with various supervisor-level address translation and protection schemes.

Supervisor mode is deliberately restricted in terms of interactions with underlying physical hardware, such as physical memory and device interrupts, to support clean virtualization. In this spirit, certain supervisor-level facilities, including requests for timer and interprocessor interrupts, are provided by implementation-specific mechanisms. In some systems, a supervisor execution environment (SEE) provides these facilities in a manner specified by a supervisor binary interface (SBI). Other systems supply these facilities directly, through some other implementation-defined mechanism.

5.1. Supervisor CSRs

A number of CSRs are provided for the supervisor.

The supervisor should only view CSR state that should be visible to a supervisor-level operating system. In particular, there is no information about the existence (or non-existence) of higher privilege levels (machine level or other) visible in the CSRs accessible by the supervisor. Many supervisor CSRs are a subset of the equivalent machine-mode CSR, and the machine-mode chapter should be read first to help understand the supervisor-level CSR descriptions.

5.1.1. Supervisor Status Register (sstatus)

The sstatus register is an SXLEN-bit read/write register formatted as shown in Figure 40 when SXLEN=32 and Figure 41 when SXLEN=64. The sstatus register keeps track of the processor’s current operating state.

Figure 40. Supervisor-mode status register (sstatus) when SXLEN=32.

Figure 41. Supervisor-mode status register (sstatus) when SXLEN=64.

The SPP bit indicates the privilege level at which a hart was executing before entering supervisor mode. When a trap is taken, SPP is set to 0 if the trap originated from user mode, or 1 otherwise. When an SRET instruction (see Section 3.3.2) is executed to return from the trap handler, the privilege level is set to user mode if the SPP bit is 0, or supervisor mode if the SPP bit is 1; SPP is then set to 0.

The SIE bit enables or disables all interrupts in supervisor mode. When SIE is clear, interrupts are not taken while in supervisor mode. When the hart is running in user-mode, the value in SIE is ignored, and supervisor-level interrupts are enabled. The supervisor can disable individual interrupt sources using the sie CSR.

The SPIE bit indicates whether supervisor interrupts were enabled prior to trapping into supervisor mode. When a trap is taken into supervisor mode, SPIE is set to SIE, and SIE is set to 0. When an SRET instruction is executed, SIE is set to SPIE, then SPIE is set to 1.

The sstatus register is a subset of the mstatus register.

In a straightforward implementation, reading or writing any field in sstatus is equivalent to reading or writing the homonymous field in mstatus.

5.1.1.1. Base ISA Control in sstatus Register

The UXL field controls the value of XLEN for U-mode, termed UXLEN, which may differ from the value of XLEN for S-mode, termed SXLEN. The encoding of UXL is the same as that of the MXL field of misa, shown in Table 9.

When SXLEN=32, the UXL field does not exist, and UXLEN=32. When SXLEN=64, it is a WARL field that encodes the current value of UXLEN. In particular, an implementation may make UXL be a read-only field whose value always ensures that UXLEN=SXLEN.

If UXLEN≠SXLEN, instructions executed in the narrower mode must ignore source register operand bits above the configured XLEN, and must sign-extend results to fill the widest supported XLEN in the destination register.

If UXLEN  SXLEN, user-mode instruction-fetch addresses and load and store effective addresses are taken modulo . For example, when UXLEN=32 and SXLEN=64, user-mode memory accesses reference the lowest of the address space.

5.1.1.2. Memory Privilege in sstatus Register

The MXR (Make eXecutable Readable) bit modifies the privilege with which loads access virtual memory. When MXR=0, only loads from pages marked readable (R=1 in Figure 57) will succeed. When MXR=1, loads from pages marked either readable or executable (R=1 or X=1) will succeed. MXR has no effect when page-based virtual memory is not in effect.

The SUM (permit Supervisor User Memory access) bit modifies the privilege with which S-mode loads and stores access virtual memory. When SUM=0, S-mode memory accesses to pages that are accessible by U-mode (U=1 in Figure 57) will fault. When SUM=1, these accesses are permitted. SUM has no effect when page-based virtual memory is not in effect, nor when executing in U-mode. Note that S-mode can never execute instructions from user pages, regardless of the state of SUM.

SUM is read-only 0 if satp.MODE is read-only 0.

The SUM mechanism prevents supervisor software from inadvertently accessing user memory. Operating systems can execute the majority of code with SUM clear; the few code segments that should access user memory can temporarily set SUM. The SUM mechanism does not avail S-mode software of permission to execute instructions in user code pages. Legitimate uses cases for execution from user memory in supervisor context are rare in general and nonexistent in POSIX environments. However, bugs in supervisors that lead to arbitrary code execution are much easier to exploit if the supervisor exploit code can be stored in a user buffer at a virtual address chosen by an attacker. Some non-POSIX single address space operating systems do allow certain privileged software to partially execute in supervisor mode, while most programs run in user mode, all in a shared address space. This use case can be realized by mapping the physical code pages at multiple virtual addresses with different permissions, possibly with the assistance of the instruction page-fault handler to direct supervisor software to use the alternate mapping.

5.1.1.3. Endianness Control in sstatus Register

The UBE bit is a WARL field that controls the endianness of explicit memory accesses made from U-mode, which may differ from the endianness of memory accesses in S-mode. An implementation may make UBE be a read-only field that always specifies the same endianness as for S-mode.

UBE controls whether explicit load and store memory accesses made from U-mode are little-endian (UBE=0) or big-endian (UBE=1).

UBE has no effect on instruction fetches, which are implicit memory accesses that are always little-endian.

For implicit accesses to supervisor-level memory management data structures, such as page tables, S-mode endianness always applies and UBE is ignored.

Standard RISC-V ABIs are expected to be purely little-endian-only or big-endian-only, with no accommodation for mixing endianness. Nevertheless, endianness control has been defined so as to permit an OS of one endianness to execute user-mode programs of the opposite endianness.

5.1.2. Supervisor Trap Vector Base Address Register (stvec)

The stvec register is an SXLEN-bit read/write register that holds trap vector configuration, consisting of a vector base address (BASE) and a vector mode (MODE).

Figure 42. Supervisor trap vector base address register (stvec).

The BASE field in stvec is a  field that can hold any valid virtual or physical address, subject to the following alignment constraints: the address must be 4-byte aligned, and MODE settings other than Direct might impose additional alignment constraints on the value in the BASE field.

Table 20. Encoding of stvec MODE field.

Value	Name	Description
0 1 ≥2	Direct Vectored	All exceptions set pc to BASE. Asynchronous interrupts set pc to BASE+4×cause. Reserved

Table 20. Encoding of stvec MODE field. {#stvec-mode}

The encoding of the MODE field is shown in Table 20. When MODE=Direct, all traps into supervisor mode cause the pc to be set to the address in the BASE field. When MODE=Vectored, all synchronous exceptions into supervisor mode cause the pc to be set to the address in the BASE field, whereas interrupts cause the pc to be set to the address in the BASE field plus four times the interrupt cause number. For example, a supervisor-mode timer interrupt (see Table 21) causes the pc to be set to BASE+0x14. Setting MODE=Vectored may impose a stricter alignment constraint on BASE.

5.1.3. Supervisor Interrupt Registers (sip and sie)

The sip register is an SXLEN-bit read/write register containing information on pending interrupts, while sie is the corresponding SXLEN-bit read/write register containing interrupt enable bits. Interrupt cause number i (as reported in CSR scause, Section 5.1.8) corresponds with bit i in both sip and sie. Bits 15:0 are allocated to standard interrupt causes only, while bits 16 and above are designated for platform or custom use.

Figure 43. Supervisor interrupt-pending register (sip).

Figure 44. Supervisor interrupt-enable register (sie).

An interrupt i will trap to S-mode if both of the following are true: (a) either the current privilege mode is S and the SIE bit in the sstatus register is set, or the current privilege mode has less privilege than S-mode; and (b) bit i is set in both sip and sie.

These conditions for an interrupt trap to occur must be evaluated in a bounded amount of time from when an interrupt becomes, or ceases to be, pending in sip, and must also be evaluated immediately following the execution of an SRET instruction or an explicit write to a CSR on which these interrupt trap conditions expressly depend (including sip, sie and sstatus).

Interrupts to S-mode take priority over any interrupts to lower privilege modes.

Each individual bit in register sip may be writable or may be read-only. When bit i in sip is writable, a pending interrupt i can be cleared by writing 0 to this bit. If interrupt i can become pending but bit i in sip is read-only, the implementation must provide some other mechanism for clearing the pending interrupt (which may involve a call to the execution environment).

A bit in sie must be writable if the corresponding interrupt can ever become pending. Bits of sie that are not writable are read-only zero.

The standard portions (bits 15:0) of registers sip and sie are formatted as shown in Figures Figure 45 and Figure 46 respectively.

Figure 45. Standard portion (bits 15:0) of sip.

Figure 46. Standard portion (bits 15:0) of sie.

Bits sip.SEIP and sie.SEIE are the interrupt-pending and interrupt-enable bits for supervisor-level external interrupts. If implemented, SEIP is read-only in sip, and is set and cleared by the execution environment, typically through a platform-specific interrupt controller.

Bits sip.STIP and sie.STIE are the interrupt-pending and interrupt-enable bits for supervisor-level timer interrupts. If implemented, STIP is read-only in sip, and is set and cleared by the execution environment.

Bits sip.SSIP and sie.SSIE are the interrupt-pending and interrupt-enable bits for supervisor-level software interrupts. If implemented, SSIP is writable in sip and may also be set to 1 by a platform-specific interrupt controller.

Interprocessor interrupts are sent to other harts by implementation-specific means, which will ultimately cause the SSIP bit to be set in the recipient hart’s sip register.

Each standard interrupt type (SEI, STI, or SSI) may not be implemented, in which case the corresponding interrupt-pending and interrupt-enable bits are read-only zeros. All bits in sip and sie are WARL fields. The implemented interrupts may be found by writing one to every bit location in sie, then reading back to see which bit positions hold a one.

The sip and sie registers are subsets of the mip and mie registers. Reading any implemented field, or writing any writable field, of sip/sie effects a read or write of the homonymous field of mip/mie. Bits 3, 7, and 11 of sip and sie correspond to the machine-mode software, timer, and external interrupts, respectively. Since most platforms will choose not to make these interrupts delegatable from M-mode to S-mode, they are shown as 0 in Figure 45 and Figure 46.

Multiple simultaneous interrupts destined for supervisor mode are handled in the following decreasing priority order: SEI, SSI, STI.

5.1.4. Supervisor Timers and Performance Counters

Supervisor software uses the same hardware performance monitoring facility as user-mode software, including the time, cycle, and instret CSRs. The implementation should provide a mechanism to modify the counter values.

The implementation must provide a facility for scheduling timer interrupts in terms of the real-time counter, time.

5.1.5. Counter-Enable Register (scounteren)

Figure 47. Counter-enable register (scounteren)

The counter-enable register scounteren is a 32-bit register that controls the availability of the hardware performance monitoring counters to U-mode.

When the CY, TM, IR, or HPMn bit in the scounteren register is clear, attempts to read the cycle, time, instret, or hpmcountern register while executing in U-mode will cause an illegal-instruction exception. When one of these bits is set, access to the corresponding register is permitted.

scounteren must be implemented. However, any of the bits may be read-only zero, indicating reads to the corresponding counter will cause an exception when executing in U-mode. Hence, they are effectively WARL fields.

The setting of a bit in mcounteren does not affect whether the corresponding bit in scounteren is writable. However, U-mode may only access a counter if the corresponding bits in scounteren and mcounteren are both set.

5.1.6. Supervisor Scratch Register (sscratch)

The sscratch register is an SXLEN-bit read/write register, dedicated for use by the supervisor. Typically, sscratch is used to hold a pointer to the hart-local supervisor context while the hart is executing user code. At the beginning of a trap handler, sscratch is swapped with a user register to provide an initial working register.

Figure 48. Supervisor Scratch Register

5.1.7. Supervisor Exception Program Counter (sepc)

sepc is an SXLEN-bit read/write register formatted as shown in Figure 49. The low bit of sepc (sepc[0]) is always zero. On implementations that support only IALIGN=32, the two low bits (sepc[1:0]) are always zero.

If an implementation allows IALIGN to be either 16 or 32 (by changing CSR misa, for example), then, whenever IALIGN=32, bit sepc[1] is masked on reads so that it appears to be 0. This masking occurs also for the implicit read by the SRET instruction. Though masked, sepc[1] remains writable when IALIGN=32.

sepc is a WARL register that must be able to hold all valid virtual addresses. It need not be capable of holding all possible invalid addresses. Prior to writing sepc, implementations may convert an invalid address into some other invalid address that sepc is capable of holding.

When a trap is taken into S-mode, sepc is written with the virtual address of the instruction that was interrupted or that encountered the exception. Otherwise, sepc is never written by the implementation, though it may be explicitly written by software.

Figure 49. Supervisor exception program counter register.

5.1.8. Supervisor Cause Register (scause)

The scause register is an SXLEN-bit read-write register formatted as shown in Figure 50. When a trap is taken into S-mode, scause is written with a code indicating the event that caused the trap. Otherwise, scause is never written by the implementation, though it may be explicitly written by software.

The Interrupt bit in the scause register is set if the trap was caused by an interrupt. The Exception Code field contains a code identifying the last exception or interrupt. Table 21 lists the possible exception codes for the current supervisor ISAs. The Exception Code is a WLRL field. It is required to hold the values 0–31 (i.e., bits 4–0 must be implemented), but otherwise it is only guaranteed to hold supported exception codes.

Figure 50. Supervisor Cause register scause.

Table 21. Supervisor cause register (scause) values after trap. Synchronous exception priorities are given by Table 15.

Interrupt	Exception Code	Description
1 1 1 1 1 1 1 1	0 1 2-4 5 6-8 9 10-15 ≥16	Reserved Supervisor software interrupt Reserved Supervisor timer interrupt Reserved Supervisor external interrupt Reserved Designated for platform use
0 0 0 0 0 0 0 0 0 0 0 0 0 0 0 0 0 0 0 0	0 1 2 3 4 5 6 7 8 9 10-11 12 13 14 15 16-23 24-31 32-47 48-63 ≥64	Instruction address misaligned Instruction access fault Illegal instruction Breakpoint Load address misaligned Load access fault Store/AMO address misaligned Store/AMO access fault Environment call from U-mode Environment call from S-mode Reserved Instruction page fault Load page fault Reserved Store/AMO page fault Reserved Designated for custom use Reserved Designated for custom use Reserved

Table 21. Supervisor cause register (scause) values after trap. Synchronous exception priorities are given by Table 15. {#scauses}

5.1.9. Supervisor Trap Value (stval) Register

The stval register is an SXLEN-bit read-write register formatted as shown in Figure 51. When a trap is taken into S-mode, stval is written with exception-specific information to assist software in handling the trap. Otherwise, stval is never written by the implementation, though it may be explicitly written by software. The hardware platform will specify which exceptions must set stval informatively and which may unconditionally set it to zero.

If stval is written with a nonzero value when a breakpoint, address-misaligned, access-fault, or page-fault exception occurs on an instruction fetch, load, or store, then stval will contain the faulting virtual address.

Figure 51. Supervisor Trap Value register.

If stval is written with a nonzero value when a misaligned load or store causes an access-fault or page-fault exception, then stval will contain the virtual address of the portion of the access that caused the fault.

If stval is written with a nonzero value when an instruction access-fault or page-fault exception occurs on a system with variable-length instructions, then stval will contain the virtual address of the portion of the instruction that caused the fault, while sepc will point to the beginning of the instruction.

The stval register can optionally also be used to return the faulting instruction bits on an illegal-instruction exception (sepc points to the faulting instruction in memory). If stval is written with a nonzero value when an illegal-instruction exception occurs, then stval will contain the shortest of:

• the actual faulting instruction

• the first ILEN bits of the faulting instruction

• the first SXLEN bits of the faulting instruction

The value loaded into stval on an illegal-instruction exception is right-justified and all unused upper bits are cleared to zero.

For other traps, stval is set to zero, but a future standard may redefine stval’s setting for other traps.

stval is a WARL register that must be able to hold all valid virtual addresses and the value 0. It need not be capable of holding all possible invalid addresses. Prior to writing stval, implementations may convert an invalid address into some other invalid address that stval is capable of holding. If the feature to return the faulting instruction bits is implemented, stval must also be able to hold all values less than , where is the smaller of SXLEN and ILEN.

5.1.10. Supervisor Environment Configuration Register (senvcfg)

The senvcfg CSR is an SXLEN-bit read/write register, formatted as shown in Figure 52, that controls certain characteristics of the U-mode execution environment.

Figure 52. Supervisor environment configuration register (senvcfg).

If bit FIOM (Fence of I/O implies Memory) is set to one in senvcfg, FENCE instructions executed in U-mode are modified so the requirement to order accesses to device I/O implies also the requirement to order main memory accesses. Table 22 details the modified interpretation of FENCE instruction bits PI, PO, SI, and SO in U-mode when FIOM=1.

Similarly, for U-mode when FIOM=1, if an atomic instruction that accesses a region ordered as device I/O has its aq and/or rl bit set, then that instruction is ordered as though it accesses both device I/O and memory.

If satp.MODE is read-only zero (always Bare), the implementation may make FIOM read-only zero.

Table 22. Modified interpretation of FENCE predecessor and successor sets in U-mode when FIOM=1.

Instruction bit	Meaning when set
PI PO	Predecessor device input and memory reads (PR implied) Predecessor device output and memory writes (PW implied)
SI SO	Successor device input and memory reads (SR implied) Successor device output and memory writes (SW implied)

Table 22. Modified interpretation of FENCE predecessor and successor sets in U-mode when FIOM=1. {#senvcfg-FIOM}

Bit FIOM exists for a specific circumstance when an I/O device is being emulated for U-mode and both of the following are true: (a) the emulated device has a memory buffer that should be I/O space but is actually mapped to main memory via address translation, and (b) multiple physical harts are involved in accessing this emulated device from U-mode. A hypervisor running in S-mode without the benefit of the hypervisor extension of Chapter 9 may need to emulate a device for U-mode if paravirtualization cannot be employed. If the same hypervisor provides a virtual machine (VM) with multiple virtual harts, mapped one-to-one to real harts, then multiple harts may concurrently access the emulated device, perhaps because: (a) the guest OS within the VM assigns device interrupt handling to one hart while the device is also accessed by a different hart outside of an interrupt handler, or (b) control of the device (or partial control) is being migrated from one hart to another, such as for interrupt load balancing within the VM. For such cases, guest software within the VM is expected to properly coordinate access to the (emulated) device across multiple harts using mutex locks and/or interprocessor interrupts as usual, which in part entails executing I/O fences. But those I/O fences may not be sufficient if some of the device ``I/O'' is actually main memory, unknown to the guest. Setting FIOM=1 modifies those fences (and all other I/O fences executed in U-mode) to include main memory, too. Software can always avoid the need to set FIOM by never using main memory to emulate a device memory buffer that should be I/O space. However, this choice usually requires trapping all U-mode accesses to the emulated buffer, which might have a noticeable impact on performance. The alternative offered by FIOM is sufficiently inexpensive to implement that we consider it worth supporting even if only rarely enabled.

The definition of the CBZE field will be furnished by the forthcoming Zicboz extension. Its allocation within senvcfg may change prior to the ratification of that extension.

The definitions of the CBCFE and CBIE fields will be furnished by the forthcoming Zicbom extension. Their allocations within senvcfg may change prior to the ratification of that extension.

5.1.11. Supervisor Address Translation and Protection (satp) Register

The satp register is an SXLEN-bit read/write register, formatted as shown in Figure 53 for SXLEN=32 and Figure 54 for SXLEN=64, which controls supervisor-mode address translation and protection. This register holds the physical page number (PPN) of the root page table, i.e., its supervisor physical address divided by ; an address space identifier (ASID), which facilitates address-translation fences on a per-address-space basis; and the MODE field, which selects the current address-translation scheme. Further details on the access to this register are described in Section 3.1.6.5.

Figure 53. Supervisor address translation and protection register satp when SXLEN=32.

Storing a PPN in satp, rather than a physical address, supports a physical address space larger than 4 GiB for RV32. The satp.PPN field might not be capable of holding all physical page numbers. Some platform standards might place constraints on the values satp.PPN may assume, e.g., by requiring that all physical page numbers corresponding to main memory be representable.

Figure 54. Supervisor address translation and protection register satp when SXLEN=64, for MODE values Bare, Sv39, Sv38, and Sv57.

We store the ASID and the page table base address in the same CSR to allow the pair to be changed atomically on a context switch. Swapping them non-atomically could pollute the old virtual address space with new translations, or vice-versa. This approach also slightly reduces the cost of a context switch.

Table 23 shows the encodings of the MODE field when SXLEN=32 and SXLEN=64. When MODE=Bare, supervisor virtual addresses are equal to supervisor physical addresses, and there is no additional memory protection beyond the physical memory protection scheme described in Section 3.7. To select MODE=Bare, software must write zero to the remaining fields of satp (bits 30–0 when SXLEN=32, or bits 59–0 when SXLEN=64). Attempting to select MODE=Bare with a nonzero pattern in the remaining fields has an UNSPECIFIED effect on the value that the remaining fields assume and an UNSPECIFIED effect on address translation and protection behavior.

When SXLEN=32, the satp encodings corresponding to MODE=Bare and ASID[8:7]=3 are designated for custom use, whereas the encodings corresponding to MODE=Bare and ASID[8:7]≠3 are reserved for future standard use. When SXLEN=64, all satp encodings corresponding to MODE=Bare are reserved for future standard use.

Version 1.11 of this standard stated that the remaining fields in satp had no effect when MODE=Bare. Making these fields reserved facilitates future definition of additional translation and protection modes, particularly in RV32, for which all patterns of the existing MODE field have already been allocated.

When SXLEN=32, the only other valid setting for MODE is Sv32, a paged virtual-memory scheme described in Section 5.3.

When SXLEN=64, three paged virtual-memory schemes are defined: Sv39, Sv48, and Sv57, described in Section 5.4, Section 5.5, and Section 5.6, respectively. One additional scheme, Sv64, will be defined in a later version of this specification. The remaining MODE settings are reserved for future use and may define different interpretations of the other fields in satp.

Implementations are not required to support all MODE settings, and if satp is written with an unsupported MODE, the entire write has no effect; no fields in satp are modified.

The number of ASID bits is  and may be zero. The number of implemented ASID bits, termed ASIDLEN, may be determined by writing one to every bit position in the ASID field, then reading back the value in satp to see which bit positions in the ASID field hold a one. The least-significant bits of ASID are implemented first: that is, if ASIDLEN  0, ASID[ASIDLEN-1:0] is writable. The maximal value of ASIDLEN, termed ASIDMAX, is 9 for Sv32 or 16 for Sv39, Sv48, and Sv57.

Table 23. Encoding of satp MODE field.

SXLEN=32
Value	Name	Description
0 1	Bare Sv32	No translation or protection. Page-based 32-bit virtual addressing (see Section 5.3).
SXLEN=64
Value	Name	Description
0 1-7 8 9 10 11 12-13 14-15	Bare - Sv39 Sv48 Sv57 Sv64 - -	No translation or protection. Reserved for standard use Page-based 39-bit virtual addressing (see Section 5.4). Page-based 48-bit virtual addressing (see Section 5.5). Page-based 57-bit virtual addressing (see Section 5.6). Reserved for page-based 64-bit virtual addressing. Reserved for standard use Designated for custom use

Table 23. Encoding of satp MODE field. {#satp-mode}

For many applications, the choice of page size has a substantial performance impact. A large page size increases TLB reach and loosens the associativity constraints on virtually indexed, physically tagged caches. At the same time, large pages exacerbate internal fragmentation, wasting physical memory and possibly cache capacity. After much deliberation, we have settled on a conventional page size of 4 KiB for both RV32 and RV64. We expect this decision to ease the porting of low-level runtime software and device drivers. The TLB reach problem is ameliorated by transparent superpage support in modern operating systems. (Navarro et al., 2002) Additionally, multi-level TLB hierarchies are quite inexpensive relative to the multi-level cache hierarchies whose address space they map.

The satp register is considered active when the effective privilege mode is S-mode or U-mode. Executions of the address-translation algorithm may only begin using a given value of satp when satp is active.

Translations that began while satp was active are not required to complete or terminate when satp is no longer active, unless an SFENCE.VMA instruction matching the address and ASID is executed. The SFENCE.VMA instruction must be used to ensure that updates to the address-translation data structures are observed by subsequent implicit reads to those structures by a hart.

Note that writing satp does not imply any ordering constraints between page-table updates and subsequent address translations, nor does it imply any invalidation of address-translation caches. If the new address space’s page tables have been modified, or if an ASID is reused, it may be necessary to execute an SFENCE.VMA instruction (see Section 5.2.1) after, or in some cases before, writing satp.

Not imposing upon implementations to flush address-translation caches upon satp writes reduces the cost of context switches, provided a sufficiently large ASID space.

5.2. Supervisor Instructions

In addition to the SRET instruction defined in Section 3.3.2, one new supervisor-level instruction is provided.

5.2.1. Supervisor Memory-Management Fence Instruction

The supervisor memory-management fence instruction SFENCE.VMA is used to synchronize updates to in-memory memory-management data structures with current execution. Instruction execution causes implicit reads and writes to these data structures; however, these implicit references are ordinarily not ordered with respect to explicit loads and stores. Executing an SFENCE.VMA instruction guarantees that any previous stores already visible to the current RISC-V hart are ordered before certain implicit references by subsequent instructions in that hart to the memory-management data structures. The specific set of operations ordered by SFENCE.VMA is determined by rs1 and rs2, as described below. SFENCE.VMA is also used to invalidate entries in the address-translation cache associated with a hart (see Section 5.3.2). Further details on the behavior of this instruction are described in Section 3.1.6.5 and Section 3.7.2.

The SFENCE.VMA is used to flush any local hardware caches related to address translation. It is specified as a fence rather than a TLB flush to provide cleaner semantics with respect to which instructions are affected by the flush operation and to support a wider variety of dynamic caching structures and memory-management schemes. SFENCE.VMA is also used by higher privilege levels to synchronize page table writes and the address translation hardware.

SFENCE.VMA orders only the local hart’s implicit references to the memory-management data structures.

Consequently, other harts must be notified separately when the memory-management data structures have been modified. One approach is to use 1) a local data fence to ensure local writes are visible globally, then 2) an interprocessor interrupt to the other thread, then 3) a local SFENCE.VMA in the interrupt handler of the remote thread, and finally 4) signal back to originating thread that operation is complete. This is, of course, the RISC-V analog to a TLB shootdown.

For the common case that the translation data structures have only been modified for a single address mapping (i.e., one page or superpage), rs1 can specify a virtual address within that mapping to effect a translation fence for that mapping only. Furthermore, for the common case that the translation data structures have only been modified for a single address-space identifier, rs2 can specify the address space. The behavior of SFENCE.VMA depends on rs1 and rs2 as follows:

• If rs1=x0 and rs2=x0, the fence orders all reads and writes made to any level of the page tables, for all address spaces. The fence also invalidates all address-translation cache entries, for all address spaces.

• If rs1=x0 and rs2≠x0, the fence orders all reads and writes made to any level of the page tables, but only for the address space identified by integer register rs2. Accesses to global mappings (see Section 5.3.1) are not ordered. The fence also invalidates all address-translation cache entries matching the address space identified by integer register rs2, except for entries containing global mappings.

• If rs1≠x0 and rs2=x0, the fence orders only reads and writes made to leaf page table entries corresponding to the virtual address in rs1, for all address spaces. The fence also invalidates all address-translation cache entries that contain leaf page table entries corresponding to the virtual address in rs1, for all address spaces.

• If rs1≠x0 and rs2≠x0, the fence orders only reads and writes made to leaf page table entries corresponding to the virtual address in rs1, for the address space identified by integer register rs2. Accesses to global mappings are not ordered. The fence also invalidates all address-translation cache entries that contain leaf page table entries corresponding to the virtual address in rs1 and that match the address space identified by integer register rs2, except for entries containing global mappings.

If the value held in rs1 is not a valid virtual address, then the SFENCE.VMA instruction has no effect. No exception is raised in this case.

When rs2≠x0, bits SXLEN-1:ASIDMAX of the value held in rs2 are reserved for future standard use. Until their use is defined by a standard extension, they should be zeroed by software and ignored by current implementations. Furthermore, if ASIDLEN<ASIDMAX, the implementation shall ignore bits ASIDMAX-1:ASIDLEN of the value held in rs2.

It is always legal to over-fence, e.g., by fencing only based on a subset of the bits in rs1 and/or rs2, and/or by simply treating all SFENCE.VMA instructions as having rs1=x0 and/or rs2=x0. For example, simpler implementations can ignore the virtual address in rs1 and the ASID value in rs2 and always perform a global fence. The choice not to raise an exception when an invalid virtual address is held in rs1 facilitates this type of simplification.

An implicit read of the memory-management data structures may return any translation for an address that was valid at any time since the most recent SFENCE.VMA that subsumes that address. The ordering implied by SFENCE.VMA does not place implicit reads and writes to the memory-management data structures into the global memory order in a way that interacts cleanly with the standard RVWMO ordering rules. In particular, even though an SFENCE.VMA orders prior explicit accesses before subsequent implicit accesses, and those implicit accesses are ordered before their associated explicit accesses, SFENCE.VMA does not necessarily place prior explicit accesses before subsequent explicit accesses in the global memory order. These implicit loads also need not otherwise obey normal program order semantics with respect to prior loads or stores to the same address.

A consequence of this specification is that an implementation may use any translation for an address that was valid at any time since the most recent SFENCE.VMA that subsumes that address. In particular, if a leaf PTE is modified but a subsuming SFENCE.VMA is not executed, either the old translation or the new translation will be used, but the choice is unpredictable. The behavior is otherwise well-defined. In a conventional TLB design, it is possible for multiple entries to match a single address if, for example, a page is upgraded to a superpage without first clearing the original non-leaf PTE’s valid bit and executing an SFENCE.VMA with rs1=x0. In this case, a similar remark applies: it is unpredictable whether the old non-leaf PTE or the new leaf PTE is used, but the behavior is otherwise well defined. Another consequence of this specification is that it is generally unsafe to update a PTE using a set of stores of a width less than the width of the PTE, as it is legal for the implementation to read the PTE at any time, including when only some of the partial stores have taken effect. This specification permits the caching of PTEs whose V (Valid) bit is clear. Operating systems must be written to cope with this possibility, but implementers are reminded that eagerly caching invalid PTEs will reduce performance by causing additional page faults.

Implementations must only perform implicit reads of the translation data structures pointed to by the current contents of the satp register or a subsequent valid (V=1) translation data structure entry, and must only raise exceptions for implicit accesses that are generated as a result of instruction execution, not those that are performed speculatively.

Changes to the sstatus fields SUM and MXR take effect immediately, without the need to execute an SFENCE.VMA instruction. Changing satp.MODE from Bare to other modes and vice versa also takes effect immediately, without the need to execute an SFENCE.VMA instruction. Likewise, changes to satp.ASID take effect immediately.

The following common situations typically require executing an SFENCE.VMA instruction: When software recycles an ASID (i.e., reassociates it with a different page table), it should first change satp to point to the new page table using the recycled ASID, then execute SFENCE.VMA with rs1=x0 and rs2 set to the recycled ASID. Alternatively, software can execute the same SFENCE.VMA instruction while a different ASID is loaded into satp, provided the next time satp is loaded with the recycled ASID, it is simultaneously loaded with the new page table. If the implementation does not provide ASIDs, or software chooses to always use ASID 0, then after every satp write, software should execute SFENCE.VMA with rs1=x0. In the common case that no global translations have been modified, rs2 should be set to a register other than x0 but which contains the value zero, so that global translations are not flushed. If software modifies a non-leaf PTE, it should execute SFENCE.VMA with rs1=x0. If any PTE along the traversal path had its G bit set, rs2 must be x0; otherwise, rs2 should be set to the ASID for which the translation is being modified. If software modifies a leaf PTE, it should execute SFENCE.VMA with rs1 set to a virtual address within the page. If any PTE along the traversal path had its G bit set, rs2 must be x0; otherwise, rs2 should be set to the ASID for which the translation is being modified. For the special cases of increasing the permissions on a leaf PTE and changing an invalid PTE to a valid leaf, software may choose to execute the SFENCE.VMA lazily. After modifying the PTE but before executing SFENCE.VMA, either the new or old permissions will be used. In the latter case, a page-fault exception might occur, at which point software should execute SFENCE.VMA in accordance with the previous bullet point.

If a hart employs an address-translation cache, that cache must appear to be private to that hart. In particular, the meaning of an ASID is local to a hart; software may choose to use the same ASID to refer to different address spaces on different harts.

A future extension could redefine ASIDs to be global across the SEE, enabling such options as shared translation caches and hardware support for broadcast TLB shootdown. However, as OSes have evolved to significantly reduce the scope of TLB shootdowns using novel ASID-management techniques, we expect the local-ASID scheme to remain attractive for its simplicity and possibly better scalability.

For implementations that make satp.MODE read-only zero (always Bare), attempts to execute an SFENCE.VMA instruction might raise an illegal-instruction exception.

5.3. Sv32: Page-Based 32-bit Virtual-Memory Systems

When Sv32 is written to the MODE field in the satp register (see Section 5.1.11), the supervisor operates in a 32-bit paged virtual-memory system. In this mode, supervisor and user virtual addresses are translated into supervisor physical addresses by traversing a radix-tree page table. Sv32 is supported when SXLEN=32 and is designed to include mechanisms sufficient for supporting modern Unix-based operating systems.

The initial RISC-V paged virtual-memory architectures have been designed as straightforward implementations to support existing operating systems. We have architected page table layouts to support a hardware page-table walker. Software TLB refills are a performance bottleneck on high-performance systems, and are especially troublesome with decoupled specialized coprocessors. An implementation can choose to implement software TLB refills using a machine-mode trap handler as an extension to M-mode. Some ISAs architecturally expose virtually indexed, physically tagged caches, in that accesses to the same physical address via different virtual addresses might not be coherent unless the virtual addresses lie within the same cache set. Implicitly, this specification does not permit such behavior to be architecturally exposed.

5.3.1. Addressing and Memory Protection

Sv32 implementations support a 32-bit virtual address space, divided into pages. An Sv32 virtual address is partitioned into a virtual page number (VPN) and page offset, as shown in Figure 55. When Sv32 virtual memory mode is selected in the MODE field of the satp register, supervisor virtual addresses are translated into supervisor physical addresses via a two-level page table. The 20-bit VPN is translated into a 22-bit physical page number (PPN), while the 12-bit page offset is untranslated. The resulting supervisor-level physical addresses are then checked using any physical memory protection structures (Section 3.7), before being directly converted to machine-level physical addresses. If necessary, supervisor-level physical addresses are zero-extended to the number of physical address bits found in the implementation.

For example, consider an RV32 system supporting 34 bits of physical address. When the value of satp.MODE is Sv32, a 34-bit physical address is produced directly, and therefore no zero-extension is needed. When the value of satp.MODE is Bare, the 32-bit virtual address is translated (unmodified) into a 32-bit physical address, and then that physical address is zero-extended into a 34-bit machine-level physical address.

Figure 55. Sv32 virtual address.

Sv32 page tables consist of 2¹⁰ page-table entries (PTEs), each of four bytes. A page table is exactly the size of a page and must always be aligned to a page boundary. The physical page number of the root page table is stored in the satp register.

Figure 56. SV32 physical address.

Figure 57. Sv32 page table entry.

The PTE format for Sv32 is shown in Figure 57. The V bit indicates whether the PTE is valid; if it is 0, all other bits in the PTE are don’t-cares and may be used freely by software. The permission bits, R, W, and X, indicate whether the page is readable, writable, and executable, respectively. When all three are zero, the PTE is a pointer to the next level of the page table; otherwise, it is a leaf PTE. Writable pages must also be marked readable; the contrary combinations are reserved for future use. Table 24 summarizes the encoding of the permission bits.

Table 24. Encoding of PTE R/W/X fields.

X	W	R	Meaning
0 0 0 0 1 1 1 1	0 0 1 1 0 0 1 1	0 1 0 1 0 1 0 1	Pointer to next level of page table. Read-only page. Reserved for future use. Read-write page. Execute-only page. Read-execute page. Reserved for future use. Read-write-execute page.

Table 24. Encoding of PTE R/W/X fields. {#pteperm}

Attempting to fetch an instruction from a page that does not have execute permissions raises a fetch page-fault exception. Attempting to execute a load or load-reserved instruction whose effective address lies within a page without read permissions raises a load page-fault exception. Attempting to execute a store, store-conditional, or AMO instruction whose effective address lies within a page without write permissions raises a store page-fault exception.

AMOs never raise load page-fault exceptions. Since any unreadable page is also unwritable, attempting to perform an AMO on an unreadable page always raises a store page-fault exception.

The U bit indicates whether the page is accessible to user mode. U-mode software may only access the page when U=1. If the SUM bit in the sstatus register is set, supervisor mode software may also access pages with U=1. However, supervisor code normally operates with the SUM bit clear, in which case, supervisor code will fault on accesses to user-mode pages. Irrespective of SUM, the supervisor may not execute code on pages with U=1.

An alternative PTE format would support different permissions for supervisor and user. We omitted this feature because it would be largely redundant with the SUM mechanism (see Section 5.1.1.2) and would require more encoding space in the PTE.

The G bit designates a global mapping. Global mappings are those that exist in all address spaces. For non-leaf PTEs, the global setting implies that all mappings in the subsequent levels of the page table are global. Note that failing to mark a global mapping as global merely reduces performance, whereas marking a non-global mapping as global is a software bug that, after switching to an address space with a different non-global mapping for that address range, can unpredictably result in either mapping being used.

Global mappings need not be stored redundantly in address-translation caches for multiple ASIDs. Additionally, they need not be flushed from local address-translation caches when an SFENCE.VMA instruction is executed with rs2≠x0.

The RSW field is reserved for use by supervisor software; the implementation shall ignore this field.

Each leaf PTE contains an accessed (A) and dirty (D) bit. The A bit indicates the virtual page has been read, written, or fetched from since the last time the A bit was cleared. The D bit indicates the virtual page has been written since the last time the D bit was cleared.

Two schemes to manage the A and D bits are permitted:

• When a virtual page is accessed and the A bit is clear, or is written and the D bit is clear, a page-fault exception is raised.

• When a virtual page is accessed and the A bit is clear, or is written and the D bit is clear, the implementation sets the corresponding bit(s) in the PTE. The PTE update must be atomic with respect to other accesses to the PTE, and must atomically check that the PTE is valid and grants sufficient permissions. Updates of the A bit may be performed as a result of speculation, but updates to the D bit must be exact (i.e., not speculative), and observed in program order by the local hart. Furthermore, the PTE update must appear in the global memory order no later than the explicit memory access, or any subsequent explicit memory access to that virtual page by the local hart. The ordering on loads and stores provided by FENCE instructions and the acquire/release bits on atomic instructions also orders the PTE updates associated with those loads and stores as observed by remote harts.

  The PTE update is not required to be atomic with respect to the explicit memory access that caused the update, and the sequence is interruptible. However, the hart must not perform the explicit memory access before the PTE update is globally visible.

All harts in a system must employ the same PTE-update scheme as each other.

Prior versions of this specification required PTE A bit updates to be exact, but allowing the A bit to be updated as a result of speculation simplifies the implementation of address translation prefetchers. System software typically uses the A bit as a page replacement policy hint, but does not require exactness for functional correctness. On the other hand, D bit updates are still required to be exact and performed in program order, as the D bit affects the functional correctness of page eviction. Implementations are of course still permitted to perform both A and D bit updates only in an exact manner. In both cases, requiring atomicity ensures that the PTE update will not be interrupted by other intervening writes to the page table, as such interruptions could lead to A/D bits being set on PTEs that have been reused for other purposes, on memory that has been reclaimed for other purposes, and so on. Simple implementations may instead generate page-fault exceptions. The A and D bits are never cleared by the implementation. If the supervisor software does not rely on accessed and/or dirty bits, e.g. if it does not swap memory pages to secondary storage or if the pages are being used to map I/O space, it should always set them to 1 in the PTE to improve performance.

Any level of PTE may be a leaf PTE, so in addition to 4 KiB pages, Sv32 supports 4 MiB megapages. A megapage must be virtually and physically aligned to a 4 MiB boundary; a page-fault exception is raised if the physical address is insufficiently aligned.

For non-leaf PTEs, the D, A, and U bits are reserved for future standard use. Until their use is defined by a standard extension, they must be cleared by software for forward compatibility.

For implementations with both page-based virtual memory and the "A" standard extension, the LR/SC reservation set must lie completely within a single base physical page (i.e., a naturally aligned 4 KiB physical-memory region).

On some implementations, misaligned loads, stores, and instruction fetches may also be decomposed into multiple accesses, some of which may succeed before a page-fault exception occurs. In particular, a portion of a misaligned store that passes the exception check may become visible, even if another portion fails the exception check. The same behavior may manifest for stores wider than XLEN bits (e.g., the FSD instruction in RV32D), even when the store address is naturally aligned.

5.3.2. Virtual Address Translation Process

A virtual address va is translated into a physical address pa as follows:

1. Let a be satp.ppn×PAGESIZE, and let i=LEVELS-1. (For Sv32, PAGESIZE=2¹² and LEVELS=2.) The satp register must be active, i.e., the effective privilege mode must be S-mode or U-mode.

2. Let pte be the value of the PTE at address a+va.vpn[i]×PTESIZE. (For Sv32, PTESIZE=4.) If accessing pte violates a PMA or PMP check, raise an access-fault exception corresponding to the original access type.

3. If pte.v=0, or if pte.r=0 and pte.w=1, or if any bits or encodings that are reserved for future standard use are set within pte, stop and raise a page-fault exception corresponding to the original access type.

4. Otherwise, the PTE is valid. If pte.r=1 or pte.x=1, go to step 5. Otherwise, this PTE is a pointer to the next level of the page table. Let i=i-1. If i<0, stop and raise a page-fault exception corresponding to the original access type. Otherwise, let a=pte.ppn×PAGESIZE and go to step 2.

5. A leaf PTE has been found. Determine if the requested memory access is allowed by the pte.r, pte.w, pte.x, and pte.u bits, given the current privilege mode and the value of the SUM and MXR fields of the mstatus register. If not, stop and raise a page-fault exception corresponding to the original access type.

6. If i>0 and pte.ppn[i-1:0] ≠ 0, this is a misaligned superpage; stop and raise a page-fault exception corresponding to the original access type.

7. If pte.a=0, or if the original memory access is a store and pte.d=0, either raise a page-fault exception corresponding to the original access type, or:

   • If a store to pte would violate a PMA or PMP check, raise an access-fault exception corresponding to the original access type.

   • Perform the following steps atomically:

     • Compare pte to the value of the PTE at address a+va.vpn[i]×PTESIZE.

     • If the values match, set pte.a to 1 and, if the original memory access is a store, also set pte.d to 1.

     • If the comparison fails, return to step 2.

8. The translation is successful. The translated physical address is given as follows:

   • pa.pgoff = va.pgoff.

   • If i>0, then this is a superpage translation and pa.ppn[i-1:0] = va.vpn[i-1:0].

   • pa.ppn[LEVELS-1:i] = pte.ppn[LEVELS-1:i].

All implicit accesses to the address-translation data structures in this algorithm are performed using width PTESIZE.

This implies, for example, that an Sv48 implementation may not use two separate 4B reads to non-atomically access a single 8B PTE, and that A/D bit updates performed by the implementation are treated as atomically updating the entire PTE, rather than just the A and/or D bit alone (even though the PTE value does not otherwise change).

The results of implicit address-translation reads in step 2 may be held in a read-only, incoherent address-translation cache but not shared with other harts. The address-translation cache may hold an arbitrary number of entries, including an arbitrary number of entries for the same address and ASID. Entries in the address-translation cache may then satisfy subsequent step 2 reads if the ASID associated with the entry matches the ASID loaded in step 0 or if the entry is associated with a global mapping. To ensure that implicit reads observe writes to the same memory locations, an SFENCE.VMA instruction must be executed after the writes to flush the relevant cached translations.

The address-translation cache cannot be used in step 7; accessed and dirty bits may only be updated in memory directly.

It is permitted for multiple address-translation cache entries to co-exist for the same address. This represents the fact that in a conventional TLB hierarchy, it is possible for multiple entries to match a single address if, for example, a page is upgraded to a superpage without first clearing the original non-leaf PTE’s valid bit and executing an SFENCE.VMA with rs1=x0, or if multiple TLBs exist in parallel at a given level of the hierarchy. In this case, just as if an SFENCE.VMA is not executed between a write to the memory-management tables and subsequent implicit read of the same address: it is unpredictable whether the old non-leaf PTE or the new leaf PTE is used, but the behavior is otherwise well defined.

Implementations may also execute the address-translation algorithm speculatively at any time, for any virtual address, as long as satp is active (as defined in Section 5.1.11). Such speculative executions have the effect of pre-populating the address-translation cache.

Speculative executions of the address-translation algorithm behave as non-speculative executions of the algorithm do, except that they must not set the dirty bit for a PTE, they must not trigger an exception, and they must not create address-translation cache entries if those entries would have been invalidated by any SFENCE.VMA instruction executed by the hart since the speculative execution of the algorithm began.

For instance, it is illegal for both non-speculative and speculative executions of the translation algorithm to begin, read the level 2 page table, pause while the hart executes an SFENCE.VMA with rs1=rs2=x0, then resume using the now-stale level 2 PTE, as subsequent implicit reads could populate the address-translation cache with stale PTEs. In many implementations, an SFENCE.VMA instruction with rs1=x0 will therefore either terminate all previously-launched speculative executions of the address-translation algorithm (for the specified ASID, if applicable), or simply wait for them to complete (in which case any address-translation cache entries created will be invalidated by the SFENCE.VMA as appropriate). Likewise, an SFENCE.VMA instruction with rs1≠x0 generally must either ensure that previously-launched speculative executions of the address-translation algorithm (for the specified ASID, if applicable) are prevented from creating new address-translation cache entries mapping leaf PTEs, or wait for them to complete. A consequence of implementations being permitted to read the translation data structures arbitrarily early and speculatively is that at any time, all page table entries reachable by executing the algorithm may be loaded into the address-translation cache. Although it would be uncommon to place page tables in non-idempotent memory, there is no explicit prohibition against doing so. Since the algorithm may only touch page tables reachable from the root page table indicated in satp, the range of addresses that an implementation’s page table walker will touch is fully under supervisor control. The algorithm does not admit the possibility of ignoring high-order PPN bits for implementations with narrower physical addresses.

5.4. Sv39: Page-Based 39-bit Virtual-Memory System

This section describes a simple paged virtual-memory system for SXLEN=64, which supports 39-bit virtual address spaces. The design of Sv39 follows the overall scheme of Sv32, and this section details only the differences between the schemes.

We specified multiple virtual memory systems for RV64 to relieve the tension between providing a large address space and minimizing address-translation cost. For many systems, 39 bits of virtual-address space is ample, and so Sv39 suffices. Sv48 increases the virtual address space to 48 bits, but increases the physical memory capacity dedicated to page tables, the latency of page-table traversals, and the size of hardware structures that store virtual addresses. Sv57 increases the virtual address space, page table capacity requirement, and translation latency even further.

5.4.1. Addressing and Memory Protection

Sv39 implementations support a 39-bit virtual address space, divided into pages. An Sv39 address is partitioned as shown in Figure 58. Instruction fetch addresses and load and store effective addresses, which are 64 bits, must have bits 63–39 all equal to bit 38, or else a page-fault exception will occur. The 27-bit VPN is translated into a 44-bit PPN via a three-level page table, while the 12-bit page offset is untranslated.

When mapping between narrower and wider addresses, RISC-V zero-extends a narrower physical address to a wider size. The mapping between 64-bit virtual addresses and the 39-bit usable address space of Sv39 is not based on zero-extension but instead follows an entrenched convention that allows an OS to use one or a few of the most-significant bits of a full-size (64-bit) virtual address to quickly distinguish user and supervisor address regions.

Figure 58. Sv39 virtual address.

Figure 59. Sv39 physical address.

Figure 60. Sv39 page table entry.

Sv39 page tables contain 2⁹ page table entries (PTEs), eight bytes each. A page table is exactly the size of a page and must always be aligned to a page boundary. The physical page number of the root page table is stored in the satp register’s PPN field.

The PTE format for Sv39 is shown in Figure 60. Bits 9-0 have the same meaning as for Sv32. Bit 63 is reserved for use by the Svnapot extension in Chapter 6. If Svnapot is not implemented, bit 63 remains reserved and must be zeroed by software for forward compatibility, or else a page-fault exception is raised. Bits 62-61 are reserved for use by the Svpbmt extension in Chapter 7. If Svpbmt is not implemented, bits 62-61 remain reserved and must be zeroed by software for forward compatibility, or else a page-fault exception is raised. Bits 60-54 are reserved for future standard use and, until their use is defined by some standard extension, must be zeroed by software for forward compatibility. If any of these bits are set, a page-fault exception is raised.

We reserved several PTE bits for a possible extension that improves support for sparse address spaces by allowing page-table levels to be skipped, reducing memory usage and TLB refill latency. These reserved bits may also be used to facilitate research experimentation. The cost is reducing the physical address space, but is presently ample. When it no longer suffices, the reserved bits that remain unallocated could be used to expand the physical address space.

Any level of PTE may be a leaf PTE, so in addition to 4 KiB pages, Sv39 supports 2 MiB megapages and 1 GiB gigapages, each of which must be virtually and physically aligned to a boundary equal to its size. A page-fault exception is raised if the physical address is insufficiently aligned.

The algorithm for virtual-to-physical address translation is the same as in Section 5.3.2, except LEVELS equals 3 and PTESIZE equals 8.

5.5. Sv48: Page-Based 48-bit Virtual-Memory System

This section describes a simple paged virtual-memory system for SXLEN=64, which supports 48-bit virtual address spaces. Sv48 is intended for systems for which a 39-bit virtual address space is insufficient. It closely follows the design of Sv39, simply adding an additional level of page table, and so this chapter only details the differences between the two schemes.

Implementations that support Sv48 must also support Sv39.

Systems that support Sv48 can also support Sv39 at essentially no cost, and so should do so to maintain compatibility with supervisor software that assumes Sv39.

5.5.1. Addressing and Memory Protection

Sv48 implementations support a 48-bit virtual address space, divided into pages. An Sv48 address is partitioned as shown in Figure 61. Instruction fetch addresses and load and store effective addresses, which are 64 bits, must have bits 63–48 all equal to bit 47, or else a page-fault exception will occur. The 36-bit VPN is translated into a 44-bit PPN via a four-level page table, while the 12-bit page offset is untranslated.

Figure 61. Sv48 virtual address.

Figure 62. Sv48 physical address.

Figure 63. Sv48 page table entry.

The PTE format for Sv48 is shown in Figure 63. Bits 63-54 and 9-0 have the same meaning as for Sv39. Any level of PTE may be a leaf PTE, so in addition to pages, Sv48 supports megapages, gigapages, and terapages, each of which must be virtually and physically aligned to a boundary equal to its size. A page-fault exception is raised if the physical address is insufficiently aligned.

The algorithm for virtual-to-physical address translation is the same as in Section 5.3.2, except LEVELS equals 4 and PTESIZE equals 8.

5.6. Sv57: Page-Based 57-bit Virtual-Memory System

This section describes a simple paged virtual-memory system designed for RV64 systems, which supports 57-bit virtual address spaces. Sv57 is intended for systems for which a 48-bit virtual address space is insufficient. It closely follows the design of Sv48, simply adding an additional level of page table, and so this chapter only details the differences between the two schemes.

Implementations that support Sv57 must also support Sv48.

Systems that support Sv57 can also support Sv48 at essentially no cost, and so should do so to maintain compatibility with supervisor software that assumes Sv48.

5.6.1. Addressing and Memory Protection

Sv57 implementations support a 57-bit virtual address space, divided into pages. An Sv57 address is partitioned as shown in Figure 64. Instruction fetch addresses and load and store effective addresses, which are 64 bits, must have bits 63–57 all equal to bit 56, or else a page-fault exception will occur. The 45-bit VPN is translated into a 44-bit PPN via a five-level page table, while the 12-bit page offset is untranslated.

Figure 64. Sv57 virtual address.

Figure 65. Sv57 physical address.

Figure 66. Sv57 page table entry.

The PTE format for Sv57 is shown in Figure 66. Bits 63–54 and 9–0 have the same meaning as for Sv39. Any level of PTE may be a leaf PTE, so in addition to pages, Sv57 supports megapages, gigapages, terapages, and petapages, each of which must be virtually and physically aligned to a boundary equal to its size. A page-fault exception is raised if the physical address is insufficiently aligned.

The algorithm for virtual-to-physical address translation is the same as in Section 5.3.2, except LEVELS equals 5 and PTESIZE equals 8.

6. "Svnapot" Standard Extension for NAPOT Translation Contiguity, Version 1.0

In Sv39, Sv48, and Sv57, when a PTE has N=1, the PTE represents a translation that is part of a range of contiguous virtual-to-physical translations with the same values for PTE bits 5–0. Such ranges must be of a naturally aligned power-of-2 (NAPOT) granularity larger than the base page size.

The Svnapot extension depends on Sv39.

Table 25. Page table entry encodings when pte.N=1

i	pte.ppn[i]	Description	pte.napot_bits
0 0 0 0 0 ≥1	x xxxx xxx1 x xxxx xx1x x xxxx x1xx x xxxx 1000 x xxxx 0xxx x xxxx xxxx	Reserved Reserved Reserved 64 KiB contiguous region Reserved Reserved	- - - 4 - -

Table 25. Page table entry encodings when pte.N=1 {#ptenapot}

NAPOT PTEs behave identically to non-NAPOT PTEs within the address-translation algorithm in Section 5.3.2, except that:

• If the encoding in pte is valid according to Table 25, then instead of returning the original value of pte, implicit reads of a NAPOT PTE return a copy of pte in which pte.ppn[i][pte.napot_bits-1:0] is replaced by vpn[i][pte.napot_bits-1:0]. If the encoding in pte is reserved according to Table 25, then a page-fault exception must be raised.

• Implicit reads of NAPOT page table entries may create address-translation cache entries mapping a + j×PTESIZE to a copy of pte in which pte.ppn[i][pte.napot_bits-1:0] is replaced by vpn[i][pte.napot_bits-1:0], for any or all j such that j >> napot_bits = vpn[i] >> napot_bits, all for the address space identified in satp as loaded by step 1.

The motivation for a NAPOT PTE is that it can be cached in a TLB as one or more entries representing the contiguous region as if it were a single (large) page covered by a single translation. This compaction can help relieve TLB pressure in some scenarios. The encoding is designed to fit within the pre-existing Sv39, Sv48, and Sv57 PTE formats so as not to disrupt existing implementations or designs that choose not to implement the scheme. It is also designed so as not to complicate the definition of the address-translation algorithm. The address translation cache abstraction captures the behavior that would result from the creation of a single TLB entry covering the entire NAPOT region. It is also designed to be consistent with implementations that support NAPOT PTEs by splitting the NAPOT region into TLB entries covering any smaller power-of-two region sizes. For example, a 64 KiB NAPOT PTE might trigger the creation of 16 standard 4 KiB TLB entries, all with contents generated from the NAPOT PTE (even if the PTEs for the other 4 KiB regions have different contents). In typical usage scenarios, NAPOT PTEs in the same region will have the same attributes, same PPNs, and same values for bits 5-0. RSW remains reserved for supervisor software control. It is the responsibility of the OS and/or hypervisor to configure the page tables in such a way that there are no inconsistencies between NAPOT PTEs and other NAPOT or non-NAPOT PTEs that overlap the same address range. If an update needs to be made, the OS generally should first mark all of the PTEs invalid, then issue SFENCE.VMA instruction(s) covering all 4 KiB regions within the range (either via a single SFENCE.VMA with rs1=x0, or with multiple SFENCE.VMA instructions with rs1≠x0), then update the PTE(s), as described in Section 5.2.1, unless any inconsistencies are known to be benign. If any inconsistencies do exist, then the effect is the same as when SFENCE.VMA is used incorrectly: one of the translations will be chosen, but the choice is unpredictable. If an implementation chooses to use a NAPOT PTE (or cached version thereof), it might not consult the PTE directly specified by the algorithm in Section 5.3.2 at all. Therefore, the D and A bits may not be identical across all mappings of the same address range even in typical use cases The operating system must query all NAPOT aliases of a page to determine whether that page has been accessed and/or is dirty. If the OS manually sets the A and/or D bits for a page, it is recommended that the OS also set the A and/or D bits for other NAPOT aliases as appropriate in order to avoid unnecessary traps. Just as with normal PTEs, TLBs are permitted to cache NAPOT PTEs whose V (Valid) bit is clear. Depending on need, the NAPOT scheme may be extended to other intermediate page sizes and/or to other levels of the page table in the future. The encoding is designed to accommodate other NAPOT sizes should that need arise. For example: __ i pte.ppn[i] Description pte.napot_bits 0 0 0 0 0 …​ 1 1 …​ x xxxx xxx1 x xxxx xx10 x xxxx x100 x xxxx 1000 x xxx1 0000 …​ x xxxx xxx1 x xxxx xx10 …​ 8 KiB contiguous region 16 KiB contiguous region 32 KiB contiguous region 64 KiB contiguous region 128 KiB contiguous region …​ 4 MiB contiguous region 8 MiB contiguous region …​ 1 2 3 4 5 …​ 1 2 …​ In such a case, an implementation may or may not support all options. The discoverability mechanism for this extension would be extended to allow system software to determine which sizes are supported. Other sizes may remain deliberately excluded, so that PPN bits not being used to indicate a valid NAPOT region size (e.g., the least-significant bit of pte.ppn[i]) may be repurposed for other uses in the future. However, in case finer-grained intermediate page size support proves not to be useful, we have chosen to standardize only 64 KiB support as a first step.

7. "Svpbmt" Standard Extension for Page-Based Memory Types, Version 1.0

In Sv39, Sv48, and Sv57, bits 62-61 of a leaf page table entry indicate the use of page-based memory types that override the PMA(s) for the associated memory pages. The encoding for the PBMT bits is captured in Table 26.

The Svpbmt extension depends on Sv39.

Table 26. Encodings for PBMT field in Sv39, Sv48, and Sv57 PTEs. Attributes not mentioned are inherited from PMA associated with the physical address.

Mode	Value	Requested Memory Attributes
PMA NC IO -	0 1 2 3	None Non-cacheable, idempotent, weakly-ordered (RVWMO), main memory Non-cacheable, non-idempotent, strongly-ordered (I/O ordering), I/O Reserved for future standard use

Table 26. Encodings for PBMT field in Sv39, Sv48, and Sv57 PTEs. Attributes not mentioned are inherited from PMA associated with the physical address. {#pbmt}

Future extensions may provide more and/or finer-grained control over which PMAs can be overridden.

For non-leaf PTEs, bits 62-61 are reserved for future standard use. Until their use is defined by a standard extension, they must be cleared by software for forward compatibility, or else a page-fault exception is raised.

For leaf PTEs, setting bits 62-61 to the value 3 is reserved for future standard use. Until this value is defined by a standard extension, using this reserved value in a leaf PTE raises a page-fault exception.

If the underlying physical memory attribute for a page is vacant, the PBMT settings do not override that.

When PBMT settings override a main memory page into I/O or vice versa, memory accesses to such pages obey the memory ordering rules of the final effective attribute, as follows.

If the underlying physical memory attribute for a page is I/O, and the page has PBMT=NC, then accesses to that page obey RVWMO. However, accesses to such pages are considered to be both I/O and main memory accesses for the purposes of FENCE, .aq, and .rl.

If the underlying physical memory attribute for a page is main memory, and the page has PBMT=IO, then accesses to that page obey strong channel 0 I/O ordering rules. However, accesses to such pages are considered to be both I/O and main memory accesses for the purposes of FENCE, .aq, and .rl.

A device driver written to rely on I/O strong ordering rules will not operate correctly if the address range is mapped with PBMT=NC. As such, this configuration is discouraged. It will often still be useful to map physical I/O regions using PBMT=NC so that write combining and speculative accesses can be performed. Such optimizations will likely improve performance when applied with adequate care.

When Svpbmt is used with non-zero PBMT encodings, it is possible for multiple virtual aliases of the same physical page to exist simultaneously with different memory attributes. It is also possible for a U-mode or S-mode mapping through a PTE with Svpbmt enabled to observe different memory attributes for a given region of physical memory than a concurrent access to the same page performed by M-mode or when MODE=Bare. In such cases, the behaviors dictated by the attributes (including coherence, which is otherwise unaffected) may be violated.

Accessing the same location using different attributes that are both non-cacheable (e.g., NC and IO) does not cause loss of coherence, but might result in weaker memory ordering than the stricter attribute ordinarily guarantees. Executing a fence iorw, iorw instruction between such accesses suffices to prevent loss of memory ordering.

Accessing the same location using different cacheability attributes may cause loss of coherence. Executing the following sequence between such accesses prevents both loss of coherence and loss of memory ordering: fence iorw, iorw, followed by cbo.flush to an address of that location, followed by a fence iorw, iorw.

It follows that, if the same location might later be referenced using the original attributes, then this sequence must be repeated beforehand. In certain cases, a weaker sequence might suffice to prevent loss of coherence. These situations will be detailed following the forthcoming formalization of the interaction of the RVWMO memory model with the instructions in the Zicbom extension.

When two-stage address translation is enabled within the H extension, the page-based memory types are also applied in two stages. First, if hgatp.MODE is not equal to zero, non-zero G-stage PTE PBMT bits override the attributes in the PMA to produce an intermediate set of attributes. Otherwise, the PMAs serve as the intermediate attributes. Second, if vsatp.MODE is not equal to zero, non-zero VS-stage PTE PBMT bits override the intermediate attributes to produce the final set of attributes used by accesses to the page in question. Otherwise, the intermediate attributes are used as the final set of attributes.

8. "Svinval" Standard Extension for Fine-Grained Address-Translation Cache Invalidation, Version 1.0

The Svinval extension splits SFENCE.VMA, HFENCE.VVMA, and HFENCE.GVMA instructions into finer-grained invalidation and ordering operations that can be more efficiently batched or pipelined on certain classes of high-performance implementation.

The SINVAL.VMA instruction invalidates any address-translation cache entries that an SFENCE.VMA instruction with the same values of rs1 and rs2 would invalidate. However, unlike SFENCE.VMA, SINVAL.VMA instructions are only ordered with respect to SFENCE.VMA, SFENCE.W.INVAL, and SFENCE.INVAL.IR instructions as defined below.

The SFENCE.W.INVAL instruction guarantees that any previous stores already visible to the current RISC-V hart are ordered before subsequent SINVAL.VMA instructions executed by the same hart. The SFENCE.INVAL.IR instruction guarantees that any previous SINVAL.VMA instructions executed by the current hart are ordered before subsequent implicit references by that hart to the memory-management data structures.

When executed in order (but not necessarily consecutively) by a single hart, the sequence SFENCE.W.INVAL, SINVAL.VMA, and SFENCE.INVAL.IR has the same effect as a hypothetical SFENCE.VMA instruction in which:

• the values of rs1 and rs2 for the SFENCE.VMA are the same as those used in the SINVAL.VMA,

• reads and writes prior to the SFENCE.W.INVAL are considered to be those prior to the SFENCE.VMA, and

• reads and writes following the SFENCE.INVAL.IR are considered to be those subsequent to the SFENCE.VMA.

If the hypervisor extension is implemented, the Svinval extension also provides two additional instructions: HINVAL.VVMA and HINVAL.GVMA. These have the same semantics as SINVAL.VMA, except that they combine with SFENCE.W.INVAL and SFENCE.INVAL.IR to replace HFENCE.VVMA and HFENCE.GVMA, respectively, instead of SFENCE.VMA. In addition, HINVAL.GVMA uses VMIDs instead of ASIDs.

SINVAL.VMA, HINVAL.VVMA, and HINVAL.GVMA require the same permissions and raise the same exceptions as SFENCE.VMA, HFENCE.VVMA, and HFENCE.GVMA, respectively. In particular, an attempt to execute any of these instructions in U-mode always raises an illegal-instruction exception, and an attempt to execute SINVAL.VMA or HINVAL.GVMA in S-mode or HS-mode when mstatus.TVM=1 also raises an illegal-instruction exception. An attempt to execute HINVAL.VVMA or HINVAL.GVMA in VS-mode or VU-mode, or to execute SINVAL.VMA in VU-mode, raises a virtual-instruction exception. When hstatus.VTVM=1, an attempt to execute SINVAL.VMA in VS-mode also raises a virtual instruction exception.

SFENCE.W.INVAL and SFENCE.INVAL.IR instructions do not need to be trapped when mstatus.TVM=1 or when hstatus.VTVM=1, as they only have ordering effects but no visible side effects. Trapping of the SINVAL.VMA instruction is sufficient to enable emulation of the intended overall TLB maintenance functionality. In typical usage, software will invalidate a range of virtual addresses in the address-translation caches by executing an SFENCE.W.INVAL instruction, executing a series of SINVAL.VMA, HINVAL.VVMA, or HINVAL.GVMA instructions to the addresses (and optionally ASIDs or VMIDs) in question, and then executing an SFENCE.INVAL.IR instruction. High-performance implementations will be able to pipeline the address-translation cache invalidation operations, and will defer any pipeline stalls or other memory ordering enforcement until an SFENCE.W.INVAL, SFENCE.INVAL.IR, SFENCE.VMA, HFENCE.GVMA, or HFENCE.VVMA instruction is executed. Simpler implementations may implement SINVAL.VMA, HINVAL.VVMA, and HINVAL.GVMA identically to SFENCE.VMA, HFENCE.VVMA, and HFENCE.GVMA, respectively, while implementing SFENCE.W.INVAL and SFENCE.INVAL.IR instructions as no-ops.

9. Hypervisor Extension, Version 1.0

This chapter describes the RISC-V hypervisor extension, which virtualizes the supervisor-level architecture to support the efficient hosting of guest operating systems atop a type-1 or type-2 hypervisor. The hypervisor extension changes supervisor mode into hypervisor-extended supervisor mode (HS-mode, or hypervisor mode for short), where a hypervisor or a hosting-capable operating system runs. The hypervisor extension also adds another stage of address translation, from guest physical addresses to supervisor physical addresses, to virtualize the memory and memory-mapped I/O subsystems for a guest operating system. HS-mode acts the same as S-mode, but with additional instructions and CSRs that control the new stage of address translation and support hosting a guest OS in virtual S-mode (VS-mode). Regular S-mode operating systems can execute without modification either in HS-mode or as VS-mode guests.

In HS-mode, an OS or hypervisor interacts with the machine through the same SBI as an OS normally does from S-mode. An HS-mode hypervisor is expected to implement the SBI for its VS-mode guest.

The hypervisor extension depends on an "I" base integer ISA with 32 x registers (RV32I or RV64I), not RV32E or RV64E, which have only 16 x registers. CSR mtval must not be read-only zero, and standard page-based address translation must be supported, either Sv32 for RV32, or a minimum of Sv39 for RV64.

The hypervisor extension is enabled by setting bit 7 in the misa CSR, which corresponds to the letter H. RISC-V harts that implement the hypervisor extension are encouraged not to hardwire misa[7], so that the extension may be disabled.

The baseline privileged architecture is designed to simplify the use of classic virtualization techniques, where a guest OS is run at user-level, as the few privileged instructions can be easily detected and trapped. The hypervisor extension improves virtualization performance by reducing the frequency of these traps. The hypervisor extension has been designed to be efficiently emulable on platforms that do not implement the extension, by running the hypervisor in S-mode and trapping into M-mode for hypervisor CSR accesses and to maintain shadow page tables. The majority of CSR accesses for type-2 hypervisors are valid S-mode accesses so need not be trapped. Hypervisors can support nested virtualization analogously.

9.1. Privilege Modes

The current virtualization mode, denoted V, indicates whether the hart is currently executing in a guest. When V=1, the hart is either in virtual S-mode (VS-mode), or in virtual U-mode (VU-mode) atop a guest OS running in VS-mode. When V=0, the hart is either in M-mode, in HS-mode, or in U-mode atop an OS running in HS-mode. The virtualization mode also indicates whether two-stage address translation is active (V=1) or inactive (V=0). Table 27 lists the possible privilege modes of a RISC-V hart with the hypervisor extension.

Table 27. Privilege modes with the hypervisor extension.

Virtualization Mode (V)	Nominal Privilege	Abbreviation	Name	Two-Stage Translation
0 0 0	U S M	U-mode HS-mode M-mode	User mode Hypervisor-extended supervisor mode Machine mode	Off Off Off
1 1	U S	VU-mode VS-mode	Virtual user mode Virtual supervisor mode	On On

Table 27. Privilege modes with the hypervisor extension. {#HPrivModes}

For privilege modes U and VU, the nominal privilege mode is U, and for privilege modes HS and VS, the nominal privilege mode is S.

HS-mode is more privileged than VS-mode, and VS-mode is more privileged than VU-mode. VS-mode interrupts are globally disabled when executing in U-mode.

This description does not consider the possibility of U-mode or VU-mode interrupts and will be revised if an extension for user-level interrupts is adopted.

9.2. Hypervisor and Virtual Supervisor CSRs

An OS or hypervisor running in HS-mode uses the supervisor CSRs to interact with the exception, interrupt, and address-translation subsystems. Additional CSRs are provided to HS-mode, but not to VS-mode, to manage two-stage address translation and to control the behavior of a VS-mode guest: hstatus, hedeleg, hideleg, hvip, hip, hie, hgeip, hgeie, henvcfg, henvcfgh, hcounteren, htimedelta, htimedeltah, htval, htinst, and hgatp.

Furthermore, several virtual supervisor CSRs (VS CSRs) are replicas of the normal supervisor CSRs. For example, vsstatus is the VS CSR that duplicates the usual sstatus CSR.

When V=1, the VS CSRs substitute for the corresponding supervisor CSRs, taking over all functions of the usual supervisor CSRs except as specified otherwise. Instructions that normally read or modify a supervisor CSR shall instead access the corresponding VS CSR. When V=1, an attempt to read or write a VS CSR directly by its own separate CSR address causes a virtual-instruction exception. (Attempts from U-mode cause an illegal-instruction exception as usual.) The VS CSRs can be accessed as themselves only from M-mode or HS-mode.

While V=1, the normal HS-level supervisor CSRs that are replaced by VS CSRs retain their values but do not affect the behavior of the machine unless specifically documented to do so. Conversely, when V=0, the VS CSRs do not ordinarily affect the behavior of the machine other than being readable and writable by CSR instructions.

Some standard supervisor CSRs (senvcfg, scounteren, and scontext, possibly others) have no matching VS CSR. These supervisor CSRs continue to have their usual function and accessibility even when V=1, except with VS-mode and VU-mode substituting for HS-mode and U-mode. Hypervisor software is expected to manually swap the contents of these registers as needed.

Matching VS CSRs exist only for the supervisor CSRs that must be duplicated, which are mainly those that get automatically written by traps or that impact instruction execution immediately after trap entry and/or right before SRET, when software alone is unable to swap a CSR at exactly the right moment. Currently, most supervisor CSRs fall into this category, but future ones might not.

In this chapter, we use the term HSXLEN to refer to the effective XLEN when executing in HS-mode, and VSXLEN to refer to the effective XLEN when executing in VS-mode.

9.2.1. Hypervisor Status Register (hstatus)

The hstatus register is an HSXLEN-bit read/write register formatted as shown in Figure 67 when HSXLEN=32 and Figure 68 when HSXLEN=64. The hstatus register provides facilities analogous to the mstatus register for tracking and controlling the exception behavior of a VS-mode guest.

Figure 67. Hypervisor status register (hstatus) when HSLEN=32

Figure 68. Hypervisor status register (hstatus) when HSXLEN=64.

The VSXL field controls the effective XLEN for VS-mode (known as VSXLEN), which may differ from the XLEN for HS-mode (HSXLEN). When HSXLEN=32, the VSXL field does not exist, and VSXLEN=32. When HSXLEN=64, VSXL is a WARL field that is encoded the same as the MXL field of misa, shown in Table 9 on page 19. In particular, an implementation may make VSXL be a read-only field whose value always ensures that VSXLEN=HSXLEN.

If HSXLEN is changed from 32 to a wider width, and if field VSXL is not restricted to a single value, it gets the value corresponding to the widest supported width not wider than the new HSXLEN.

The hstatus fields VTSR, VTW, and VTVM are defined analogously to the mstatus fields TSR, TW, and TVM, but affect execution only in VS-mode, and cause virtual-instruction exceptions instead of illegal-instruction exceptions. When VTSR=1, an attempt in VS-mode to execute SRET raises a virtual-instruction exception. When VTW=1 (and assuming mstatus.TW=0), an attempt in VS-mode to execute WFI raises a virtual-instruction exception if the WFI does not complete within an implementation-specific, bounded time limit. An implementation may have WFI always raise a virtual-instruction exception in VS-mode when VTW=1 (and mstatus.TW=0), even if there are pending globally-disabled interrupts when the instruction is executed. When VTVM=1, an attempt in VS-mode to execute SFENCE.VMA or SINVAL.VMA or to access CSR satp raises a virtual-instruction exception.

The VGEIN (Virtual Guest External Interrupt Number) field selects a guest external interrupt source for VS-level external interrupts. VGEIN is a WLRL field that must be able to hold values between zero and the maximum guest external interrupt number (known as GEILEN), inclusive. When VGEIN=0, no guest external interrupt source is selected for VS-level external interrupts. GEILEN may be zero, in which case VGEIN may be read-only zero. Guest external interrupts are explained in Section 9.2.4, and the use of VGEIN is covered further in Section 9.2.3.

Field HU (Hypervisor in U-mode) controls whether the virtual-machine load/store instructions, HLV, HLVX, and HSV, can be used also in U-mode. When HU=1, these instructions can be executed in U-mode the same as in HS-mode. When HU=0, all hypervisor instructions cause an illegal-instruction exception in U-mode.

The HU bit allows a portion of a hypervisor to be run in U-mode for greater protection against software bugs, while still retaining access to a virtual machine’s memory.

The SPV bit (Supervisor Previous Virtualization mode) is written by the implementation whenever a trap is taken into HS-mode. Just as the SPP bit in sstatus is set to the (nominal) privilege mode at the time of the trap, the SPV bit in hstatus is set to the value of the virtualization mode V at the time of the trap. When an SRET instruction is executed when V=0, V is set to SPV.

When V=1 and a trap is taken into HS-mode, bit SPVP (Supervisor Previous Virtual Privilege) is set to the nominal privilege mode at the time of the trap, the same as sstatus.SPP. But if V=0 before a trap, SPVP is left unchanged on trap entry. SPVP controls the effective privilege of explicit memory accesses made by the virtual-machine load/store instructions, HLV, HLVX, and HSV.

Without SPVP, if instructions HLV, HLVX, and HSV looked instead to sstatus.SPP for the effective privilege of their memory accesses, then, even with HU=1, U-mode could not access virtual machine memory at VS-level, because to enter U-mode using SRET always leaves SPP=0. Unlike SPP, field SPVP is untouched by transitions back-and-forth between HS-mode and U-mode.

Field GVA (Guest Virtual Address) is written by the implementation whenever a trap is taken into HS-mode. For any trap (breakpoint, address misaligned, access fault, page fault, or guest-page fault) that writes a guest virtual address to stval, GVA is set to 1. For any other trap into HS-mode, GVA is set to 0.

For breakpoint and memory access traps that write a nonzero value to stval, GVA is redundant with field SPV (the two bits are set the same) except when the explicit memory access of an HLV, HLVX, or HSV instruction causes a fault. In that case, SPV=0 but GVA=1.

The VSBE bit is a WARL field that controls the endianness of explicit memory accesses made from VS-mode. If VSBE=0, explicit load and store memory accesses made from VS-mode are little-endian, and if VSBE=1, they are big-endian. VSBE also controls the endianness of all implicit accesses to VS-level memory management data structures, such as page tables. An implementation may make VSBE a read-only field that always specifies the same endianness as HS-mode.

9.2.2. Hypervisor Trap Delegation Registers (hedeleg and hideleg)

Registers hedeleg and hideleg are HSXLEN-bit read/write registers, formatted as shown in Figure 69 and Figure 70 respectively. By default, all traps at any privilege level are handled in M-mode, though M-mode usually uses the medeleg and mideleg CSRs to delegate some traps to HS-mode. The hedeleg and hideleg CSRs allow these traps to be further delegated to a VS-mode guest; their layout is the same as medeleg and mideleg.

Figure 69. Hypervisor exception delegation register (hedeleg).

Figure 70. Hypervisor exception delegation register (hideleg).

A synchronous trap that has been delegated to HS-mode (using medeleg) is further delegated to VS-mode if V=1 before the trap and the corresponding hedeleg bit is set. Each bit of hedeleg shall be either writable or read-only zero. Many bits of hedeleg are required specifically to be writable or zero, as enumerated in Table 28. Bit 0, corresponding to instruction address misaligned exceptions, must be writable if IALIGN=32.

Requiring that certain bits of hedeleg be writable reduces some of the burden on a hypervisor to handle variations of implementation.

An interrupt that has been delegated to HS-mode (using mideleg) is further delegated to VS-mode if the corresponding hideleg bit is set. Among bits 15:0 of hideleg, bits 10, 6, and 2 (corresponding to the standard VS-level interrupts) are writable, and bits 12, 9, 5, and 1 (corresponding to the standard S-level interrupts) are read-only zeros.

When a virtual supervisor external interrupt (code 10) is delegated to VS-mode, it is automatically translated by the machine into a supervisor external interrupt (code 9) for VS-mode, including the value written to vscause on an interrupt trap. Likewise, a virtual supervisor timer interrupt (6) is translated into a supervisor timer interrupt (5) for VS-mode, and a virtual supervisor software interrupt (2) is translated into a supervisor software interrupt (1) for VS-mode. Similar translations may or may not be done for platform or custom interrupt causes (codes 16 and above).

Table 28. Bits of hedeleg that must be writable or must be read-only zero.

Bit	Attribute	Corresponding Exception
0 1 2 3 4 5 6 7 8 9 10 11 12 13 15 20 21 22 23	(See text) Writable Writable Writable Writable Writable Writable Writable Writable Read-only 0 Read-only 0 Read-only 0 Writable Writable Writable Read-only 0 Read-only 0 Read-only 0 Read-only 0	Instruction address misaligned Instruction access fault Illegal instruction Breakpoint Load address misaligned Load access fault Store/AMO address misaligned Store/AMO access fault Environment call from U-mode or VU-mode Environment call from HS-mode Environment call from VS-mode Environment call from M-mode Instruction page fault Load page fault Store/AMO page fault Instruction guest-page fault Load guest-page fault Virtual instruction Store/AMO guest-page fault

Table 28. Bits of hedeleg that must be writable or must be read-only zero. {#hedeleg-bits}

9.2.3. Hypervisor Interrupt Registers (hvip, hip, and hie)

Register hvip is an HSXLEN-bit read/write register that a hypervisor can write to indicate virtual interrupts intended for VS-mode. Bits of hvip that are not writable are read-only zeros.

Figure 71. Hypervisor virtual-interrupt-pending register(hvip).

The standard portion (bits 15:0) of hvip is formatted as shown in Figure 72. Bits VSEIP, VSTIP, and VSSIP of hvip are writable. Setting VSEIP=1 in hvip asserts a VS-level external interrupt; setting VSTIP asserts a VS-level timer interrupt; and setting VSSIP asserts a VS-level software interrupt.

Figure 72. Standard portion (bits 15:0) of hvip.

Registers hip and hie are HSXLEN-bit read/write registers that supplement HS-level’s sip and sie respectively. The hip register indicates pending VS-level and hypervisor-specific interrupts, while hie contains enable bits for the same interrupts.

Figure 73. Hypervisor interrupt-pending register (hip).

Figure 74. Hypervisor interrupt-enable register (hie).

For each writable bit in sie, the corresponding bit shall be read-only zero in both hip and hie. Hence, the nonzero bits in sie and hie are always mutually exclusive, and likewise for sip and hip.

The active bits of hip and hie cannot be placed in HS-level’s sip and sie because doing so would make it impossible for software to emulate the hypervisor extension on platforms that do not implement it in hardware.

An interrupt i will trap to HS-mode whenever all of the following are true: (a) either the current operating mode is HS-mode and the SIE bit in the sstatus register is set, or the current operating mode has less privilege than HS-mode; (b) bit i is set in both sip and sie, or in both hip and hie; and (c) bit i is not set in hideleg.

If bit i of sie is read-only zero, the same bit in register hip may be writable or may be read-only. When bit i in hip is writable, a pending interrupt i can be cleared by writing 0 to this bit. If interrupt i can become pending in hip but bit i in hip is read-only, then either the interrupt can be cleared by clearing bit i of hvip, or the implementation must provide some other mechanism for clearing the pending interrupt (which may involve a call to the execution environment).

A bit in hie shall be writable if the corresponding interrupt can ever become pending in hip. Bits of hie that are not writable shall be read-only zero.

The standard portions (bits 15:0) of registers hip and hie are formatted as shown in Figure 75 and Figure 76 respectively.

Figure 75. Standard portion (bits 15:0) of hip.

Figure 76. Standard portion (bits 15:0) of hie.

Bits hip.SGEIP and hie.SGEIE are the interrupt-pending and interrupt-enable bits for guest external interrupts at supervisor level (HS-level). SGEIP is read-only in hip, and is 1 if and only if the bitwise logical-AND of CSRs hgeip and hgeie is nonzero in any bit. (See Section 9.2.4.)

Bits hip.VSEIP and hie.VSEIE are the interrupt-pending and interrupt-enable bits for VS-level external interrupts. VSEIP is read-only in hip, and is the logical-OR of these interrupt sources:

• bit VSEIP of hvip;

• the bit of hgeip selected by hstatus.VGEIN; and

• any other platform-specific external interrupt signal directed to VS-level.

Bits hip.VSTIP and hie.VSTIE are the interrupt-pending and interrupt-enable bits for VS-level timer interrupts. VSTIP is read-only in hip, and is the logical-OR of hvip.VSTIP and any other platform-specific timer interrupt signal directed to VS-level.

Bits hip.VSSIP and hie.VSSIE are the interrupt-pending and interrupt-enable bits for VS-level software interrupts. VSSIP in hip is an alias (writable) of the same bit in hvip.

Multiple simultaneous interrupts destined for HS-mode are handled in the following decreasing priority order: SEI, SSI, STI, SGEI, VSEI, VSSI, VSTI.

9.2.4. Hypervisor Guest External Interrupt Registers (hgeip and hgeie)

The hgeip register is an HSXLEN-bit read-only register, formatted as shown in Figure 77, that indicates pending guest external interrupts for this hart. The hgeie register is an HSXLEN-bit read/write register, formatted as shown in Figure 78, that contains enable bits for the guest external interrupts at this hart. Guest external interrupt number i corresponds with bit i in both hgeip and hgeie.

Figure 77. Hypervisor guest external interrupt-pending register (hgeip).

Figure 78. Hypervisor guest external interrupt-enable register (hgeie).

Guest external interrupts represent interrupts directed to individual virtual machines at VS-level. If a RISC-V platform supports placing a physical device under the direct control of a guest OS with minimal hypervisor intervention (known as pass-through or direct assignment between a virtual machine and the physical device), then, in such circumstance, interrupts from the device are intended for a specific virtual machine. Each bit of hgeip summarizes all pending interrupts directed to one virtual hart, as collected and reported by an interrupt controller. To distinguish specific pending interrupts from multiple devices, software must query the interrupt controller.

Support for guest external interrupts requires an interrupt controller that can collect virtual-machine-directed interrupts separately from other interrupts.

The number of bits implemented in hgeip and hgeie for guest external interrupts is UNSPECIFIED and may be zero. This number is known as GEILEN. The least-significant bits are implemented first, apart from bit 0. Hence, if GEILEN is nonzero, bits GEILEN:1 shall be writable in hgeie, and all other bit positions shall be read-only zeros in both hgeip and hgeie.

The set of guest external interrupts received and handled at one physical hart may differ from those received at other harts. Guest external interrupt number i at one physical hart is typically expected not to be the same as guest external interrupt i at any other hart. For any one physical hart, the maximum number of virtual harts that may directly receive guest external interrupts is limited by GEILEN. The maximum this number can be for any implementation is 31 for RV32 and 63 for RV64, per physical hart. A hypervisor is always free to emulate devices for any number of virtual harts without being limited by GEILEN. Only direct pass-through (direct assignment) of interrupts is affected by the GEILEN limit, and the limit is on the number of virtual harts receiving such interrupts, not the number of distinct interrupts received. The number of distinct interrupts a single virtual hart may receive is determined by the interrupt controller.

Register hgeie selects the subset of guest external interrupts that cause a supervisor-level (HS-level) guest external interrupt. The enable bits in hgeie do not affect the VS-level external interrupt signal selected from hgeip by hstatus.VGEIN.

9.2.5. Hypervisor Environment Configuration Register (henvcfg)

The henvcfg CSR is a 64-bit read/write register, formatted as shown in Figure 79, that controls certain characteristics of the execution environment when virtualization mode V=1.

Figure 79. Hypervisor environment configuration register (henvcfg).

If bit FIOM (Fence of I/O implies Memory) is set to one in henvcfg, FENCE instructions executed when V=1 are modified so the requirement to order accesses to device I/O implies also the requirement to order main memory accesses. Table 29 details the modified interpretation of FENCE instruction bits PI, PO, SI, and SO when FIOM=1 and V=1.

Similarly, when FIOM=1 and V=1, if an atomic instruction that accesses a region ordered as device I/O has its aq and/or rl bit set, then that instruction is ordered as though it accesses both device I/O and memory.

Table 29. Modified interpretation of FENCE predecessor and successor sets when FIOM=1 and virtualization mode V=1.

Instruction bit	Meaning when set
PI PO	Predecessor device input and memory reads (PR implied) Predecessor device output and memory writes (PW implied)
SI SO	Successor device input and memory reads (SR implied) Successor device output and memory writes (SW implied)

Table 29. Modified interpretation of FENCE predecessor and successor sets when FIOM=1 and virtualization mode V=1. {#henvcfg-FIOM}

The PBMTE bit controls whether the Svpbmt extension is available for use in VS-stage address translation. When PBMTE=1, Svpbmt is available for VS-stage address translation. When PBMTE=0, the implementation behaves as though Svpbmt were not implemented for VS-stage address translation. If Svpbmt is not implemented, PBMTE is read-only zero.

The definition of the STCE field will be furnished by the forthcoming Sstc extension. Its allocation within henvcfg may change prior to the

The definition of the CBZE field will be furnished by the forthcoming Zicboz extension. Its allocation within henvcfg may change prior to the ratification of that extension.

The definitions of the CBCFE and CBIE fields will be furnished by the forthcoming Zicbom extension. Their allocations within henvcfg may change prior to the ratification of that extension.

When XLEN=32, henvcfgh is a 32-bit read/write register that aliases bits 63:32 of henvcfg. Register henvcfgh does not exist when XLEN=64.

9.2.6. Hypervisor Counter-Enable Register (hcounteren)

The counter-enable register hcounteren is a 32-bit register that controls the availability of the hardware performance monitoring counters to the guest virtual machine.

Figure 80. Hypervisor counter-enable register (hcounteren).

When the CY, TM, IR, or HPM_n_ bit in the hcounteren register is clear, attempts to read the cycle, time, instret, or hpmcounter n register while V=1 will cause a virtual-instruction exception if the same bit in mcounteren is 1. When one of these bits is set, access to the corresponding register is permitted when V=1, unless prevented for some other reason. In VU-mode, a counter is not readable unless the applicable bits are set in both hcounteren and scounteren.

hcounteren must be implemented. However, any of the bits may be read-only zero, indicating reads to the corresponding counter will cause an exception when V=1. Hence, they are effectively WARL fields.

9.2.7. Hypervisor Time Delta Register (htimedelta)

The htimedelta CSR is a 64-bit read/write register that contains the delta between the value of the time CSR and the value returned in VS-mode or VU-mode. That is, reading the time CSR in VS or VU mode returns the sum of the contents of htimedelta and the actual value of time.

Because overflow is ignored when summing htimedelta and time, large values of htimedelta may be used to represent negative time offsets.

Figure 81. Hypervisor time delta register.

When XLEN=32, htimedeltah is a 32-bit read/write register that aliases bits 63:32 of htimedelta. Register htimedeltah does not exist when XLEN=64.

9.2.8. Hypervisor Trap Value Register (htval)

The htval register is an HSXLEN-bit read/write register formatted as shown in Figure 82. When a trap is taken into HS-mode, htval is written with additional exception-specific information, alongside stval, to assist software in handling the trap.

Figure 82. Hypervisor trap value register (htval).

When a guest-page-fault trap is taken into HS-mode, htval is written with either zero or the guest physical address that faulted, shifted right by 2 bits. For other traps, htval is set to zero, but a future standard or extension may redefine htval’s setting for other traps.

A guest-page fault may arise due to an implicit memory access during first-stage (VS-stage) address translation, in which case a guest physical address written to htval is that of the implicit memory access that faulted—for example, the address of a VS-level page table entry that could not be read. (The guest physical address corresponding to the original virtual address is unknown when VS-stage translation fails to complete.) Additional information is provided in CSR htinst to disambiguate such situations.

Otherwise, for misaligned loads and stores that cause guest-page faults, a nonzero guest physical address in htval corresponds to the faulting portion of the access as indicated by the virtual address in stval. For instruction guest-page faults on systems with variable-length instructions, a nonzero htval corresponds to the faulting portion of the instruction as indicated by the virtual address in stval.

A guest physical address written to htval is shifted right by 2 bits to accommodate addresses wider than the current XLEN. For RV32, the hypervisor extension permits guest physical addresses as wide as 34 bits, and htval reports bits 33:2 of the address. This shift-by-2 encoding of guest physical addresses matches the encoding of physical addresses in PMP address registers (Section 3.7) and in page table entries (Section 5.3, Section 5.4, Section 5.5, and Section 5.6). If the least-significant two bits of a faulting guest physical address are needed, these bits are ordinarily the same as the least-significant two bits of the faulting virtual address in stval. For faults due to implicit memory accesses for VS-stage address translation, the least-significant two bits are instead zeros. These cases can be distinguished using the value provided in register htinst.

htval is a WARL register that must be able to hold zero and may be capable of holding only an arbitrary subset of other 2-bit-shifted guest physical addresses, if any.

Unless it has reason to assume otherwise (such as a platform standard), software that writes a value to htval should read back from htval to confirm the stored value.

9.2.9. Hypervisor Trap Instruction Register (htinst)

The htinst register is an HSXLEN-bit read/write register formatted as shown in Figure 83. When a trap is taken into HS-mode, htinst is written with a value that, if nonzero, provides information about the instruction that trapped, to assist software in handling the trap. The values that may be written to htinst on a trap are documented in Section 9.6.3.

Figure 83. Hypervisor trap instruction register (htinst).

htinst is a WARL register that need only be able to hold the values that the implementation may automatically write to it on a trap.

9.2.10. Hypervisor Guest Address Translation and Protection Register (hgatp)

The hgatp register is an HSXLEN-bit read/write register, formatted as shown in Figure 84 for HSXLEN=32 and Figure 85 for HSXLEN=64, which controls G-stage address translation and protection, the second stage of two-stage translation for guest virtual addresses (see Section 9.5). Similar to CSR satp, this register holds the physical page number (PPN) of the guest-physical root page table; a virtual machine identifier (VMID), which facilitates address-translation fences on a per-virtual-machine basis; and the MODE field, which selects the address-translation scheme for guest physical addresses. When mstatus.TVM=1, attempts to read or write hgatp while executing in HS-mode will raise an illegal-instruction exception.

Figure 84. Hypervisor guest address translation and protection register hgatp when HSXLEN=32.

Figure 85. Hypervisor guest address translation and protection register hgatp when HSXLEN=64 for MODE values Bare, Sv39x4, and Sv57x4.

Table 30 shows the encodings of the MODE field when HSXLEN=32 and HSXLEN=64. When MODE=Bare, guest physical addresses are equal to supervisor physical addresses, and there is no further memory protection for a guest virtual machine beyond the physical memory protection scheme described in Section 3.7. In this case, the remaining fields in hgatp must be set to zeros.

When HSXLEN=32, the only other valid setting for MODE is Sv32x4, which is a modification of the usual Sv32 paged virtual-memory scheme, extended to support 34-bit guest physical addresses. When HSXLEN=64, modes Sv39x4, Sv48x4, and Sv57x4 are defined as modifications of the Sv39, Sv48, and Sv57 paged virtual-memory schemes. All of these paged virtual-memory schemes are described in Section 9.5.1.

The remaining MODE settings when HSXLEN=64 are reserved for future use and may define different interpretations of the other fields in hgatp.

Table 30. Encoding of hgatp MODE field.

HSXLEN=32
Value	Name	Description
0 1	Bare Sv32x4	No translation or protection. Page-based 34-bit virtual addressing (2-bit extension of Sv32).
HSXLEN=64
Value	Name	Description
0 1-7 8 9 10 11-15	Bare — Sv39x4 Sv48x4 Sv57x4 —	No translation or protection. Reserved Page-based 41-bit virtual addressing (2-bit extension of Sv39). Page-based 50-bit virtual addressing (2-bit extension of Sv48). Page-based 59-bit virtual addressing (2-bit extension of Sv57). Reserved

Table 30. Encoding of hgatp MODE field. {#hgatp-mode}

Implementations are not required to support all defined MODE settings when HSXLEN=64.

A write to hgatp with an unsupported MODE value is not ignored as it is for satp. Instead, the fields of hgatp are in the normal way, when so indicated.

As explained in Section 9.5.1, for the paged virtual-memory schemes (Sv32x4, Sv39x4, Sv48x4, and Sv57x4), the root page table is 16 KiB and must be aligned to a 16-KiB boundary. In these modes, the lowest two bits of the physical page number (PPN) in hgatp always read as zeros. An implementation that supports only the defined paged virtual-memory schemes and/or Bare may make PPN[1:0] read-only zero.

The number of VMID bits is UNSPECIFIED and may be zero. The number of implemented VMID bits, termed VMIDLEN, may be determined by writing one to every bit position in the VMID field, then reading back the value in hgatp to see which bit positions in the VMID field hold a one. The least-significant bits of VMID are implemented first: that is, if VMIDLEN > 0, VMID[VMIDLEN-1:0] is writable. The maximal value of VMIDLEN, termed VMIDMAX, is 7 for Sv32x4 or 14 for Sv39x4, Sv48x4, and Sv57x4.

The hgatp register is considered active for the purposes of the address-translation algorithm unless the effective privilege mode is U and hstatus.HU=0.

This definition simplifies the implementation of speculative execution of HLV, HLVX, and HSV instructions.

Note that writing hgatp does not imply any ordering constraints between page-table updates and subsequent G-stage address translations. If the new virtual machine’s guest physical page tables have been modified, or if a VMID is reused, it may be necessary to execute an HFENCE.GVMA instruction (see Section 9.3.2) before or after writing hgatp.

9.2.11. Virtual Supervisor Status Register (vsstatus)

The vsstatus register is a VSXLEN-bit read/write register that is VS-mode’s version of supervisor register sstatus, formatted as shown in Figure 86 when VSXLEN=32 and Figure 87 when VSXLEN=64. When V=1, vsstatus substitutes for the usual sstatus, so instructions that normally read or modify sstatus actually access vsstatus instead.

Figure 86. Virtual supervisor status register (vstatus) when VSXLEN=32.

Figure 87. Virtual supervisor status register (vsstatus) when VSXLEN=64.

The UXL field controls the effective XLEN for VU-mode, which may differ from the XLEN for VS-mode (VSXLEN). When VSXLEN=32, the UXL field does not exist, and VU-mode XLEN=32. When VSXLEN=64, UXL is a WARL field that is encoded the same as the MXL field of misa, shown in Table 9. In particular, an implementation may make UXL be a read-only copy of field VSXL of hstatus, forcing VU-mode XLEN=VSXLEN.

If VSXLEN is changed from 32 to a wider width, and if field UXL is not restricted to a single value, it gets the value corresponding to the widest supported width not wider than the new VSXLEN.

When V=1, both vsstatus.FS and the HS-level sstatus.FS are in effect. Attempts to execute a floating-point instruction when either field is 0 (Off) raise an illegal-instruction exception. Modifying the floating-point state when V=1 causes both fields to be set to 3 (Dirty).

For a hypervisor to benefit from the extension context status, it must have its own copy in the HS-level sstatus, maintained independently of a guest OS running in VS-mode. While a version of the extension context status obviously must exist in vsstatus for VS-mode, a hypervisor cannot rely on this version being maintained correctly, given that VS-level software can change vsstatus.FS arbitrarily. If the HS-level sstatus.FS were not independently active and maintained by the hardware in parallel with vsstatus.FS while V=1, hypervisors would always be forced to conservatively swap all floating-point state when context-switching between virtual machines.

Similarly, when V=1, both vsstatus.VS and the HS-level sstatus.VS are in effect. Attempts to execute a vector instruction when either field is 0 (Off) raise an illegal-instruction exception. Modifying the vector state when V=1 causes both fields to be set to 3 (Dirty).

Read-only fields SD and XS summarize the extension context status as it is visible to VS-mode only. For example, the value of the HS-level sstatus.FS does not affect vsstatus.SD.

An implementation may make field UBE be a read-only copy of hstatus.VSBE.

When V=0, vsstatus does not directly affect the behavior of the machine, unless a virtual-machine load/store (HLV, HLVX, or HSV) or the MPRV feature in the mstatus register is used to execute a load or store as though V=1.

9.2.12. Virtual Supervisor Interrupt Registers (vsip and vsie)

The vsip and vsie registers are VSXLEN-bit read/write registers that are VS-mode’s versions of supervisor CSRs sip and sie, formatted as shown in Figure 88 and Figure 89 respectively. When V=1, vsip and vsie substitute for the usual sip and sie, so instructions that normally read or modify sip/sie actually access vsip/vsie instead. However, interrupts directed to HS-level continue to be indicated in the HS-level sip register, not in vsip, when V=1.

Figure 88. Virtual supervisor interrupt-pending register (vsip).

Figure 89. Virtual supervisor interrupt-enable register (vsie).

The standard portions (bits 15:0) of registers vsip and vsie are formatted as shown in Figure 90 and Figure 91 respectively.

Figure 90. Standard portion (bits 15:0) of vsip.

Figure 91. Standard portion (bits 15:0) of vsie.

When bit 10 of hideleg is zero, vsip.SEIP and vsie.SEIE are read-only zeros. Else, vsip.SEIP and vsie.SEIE are aliases of hip.VSEIP and hie.VSEIE.

When bit 6 of hideleg is zero, vsip.STIP and vsie.STIE are read-only zeros. Else, vsip.STIP and vsie.STIE are aliases of hip.VSTIP and hie.VSTIE.

When bit 2 of hideleg is zero, vsip.SSIP and vsie.SSIE are read-only zeros. Else, vsip.SSIP and vsie.SSIE are aliases of hip.VSSIP and hie.VSSIE.

9.2.13. Virtual Supervisor Trap Vector Base Address Register (vstvec)

The vstvec register is a VSXLEN-bit read/write register that is VS-mode’s version of supervisor register stvec, formatted as shown in Figure 92. When V=1, vstvec substitutes for the usual stvec, so instructions that normally read or modify stvec actually access vstvec instead. When V=0, vstvec does not directly affect the behavior of the machine.

Figure 92. Virtual supervisor trap vector base address register vstvec.

9.2.14. Virtual Supervisor Scratch Register (vsscratch)

The vsscratch register is a VSXLEN-bit read/write register that is VS-mode’s version of supervisor register sscratch, formatted as shown in Figure 93. When V=1, vsscratch substitutes for the usual sscratch, so instructions that normally read or modify sscratch actually access vsscratch instead. The contents of vsscratch never directly affect the behavior of the machine.

Figure 93. Virtual supervisor scratch register vsscratch.

9.2.15. Virtual Supervisor Exception Program Counter (vsepc)

The vsepc register is a VSXLEN-bit read/write register that is VS-mode’s version of supervisor register sepc, formatted as shown in Figure 94. When V=1, vsepc substitutes for the usual sepc, so instructions that normally read or modify sepc actually access vsepc instead. When V=0, vsepc does not directly affect the behavior of the machine.

vsepc is a WARL register that must be able to hold the same set of values that sepc can hold.

Figure 94. Virtual supervisor exception program counter (vsepc).

9.2.16. Virtual Supervisor Cause Register (vscause)

The vscause register is a VSXLEN-bit read/write register that is VS-mode’s version of supervisor register scause, formatted as shown in Figure 95. When V=1, vscause substitutes for the usual scause, so instructions that normally read or modify scause actually access vscause instead. When V=0, vscause does not directly affect the behavior of the machine.

vscause is a WLRL register that must be able to hold the same set of values that scause can hold.

Figure 95. Virtual supervisor cause register (vscause).

9.2.17. Virtual Supervisor Trap Value Register (vstval)

The vstval register is a VSXLEN-bit read/write register that is VS-mode’s version of supervisor register stval, formatted as shown in Figure 96. When V=1, vstval substitutes for the usual stval, so instructions that normally read or modify stval actually access vstval instead. When V=0, vstval does not directly affect the behavior of the machine.

vstval is a WARL register that must be able to hold the same set of values that stval can hold.

Figure 96. Virtual supervisor trap value register (vstval).

9.2.18. Virtual Supervisor Address Translation and Protection Register (vsatp)

The vsatp register is a VSXLEN-bit read/write register that is VS-mode’s version of supervisor register satp, formatted as shown in Figure 97 for VSXLEN=32 and Figure 98 for VSXLEN=64. When V=1, vsatp substitutes for the usual satp, so instructions that normally read or modify satp actually access vsatp instead. vsatp controls VS-stage address translation, the first stage of two-stage translation for guest virtual addresses (see Section 9.5).

Figure 97. Virtual supervisor address translation and protection register vsatp when VSXLEN=32.

Figure 98. Virtual supervisor address translation and protection register vsatp when VSXLEN=64.

The vsatp register is considered active for the purposes of the address-translation algorithm unless the effective privilege mode is U and hstatus.HU=0. However, even when vsatp is active, VS-stage page-table entries’ A bits must not be set as a result of speculative execution, unless the effective privilege mode is VS or VU.

In particular, virtual-machine load/store (HLV, HLVX, or HSV) instructions that are misspeculatively executed must not cause VS-stage A bits to be set.

When V=0, a write to vsatp with an unsupported MODE value is either ignored as it is for satp, or the fields of vsatp are treated as WARL in the normal way. However, when V=1, a write to satp with an unsupported MODE value is ignored and no write to vsatp is effected.

When V=0, vsatp does not directly affect the behavior of the machine, unless a virtual-machine load/store (HLV, HLVX, or HSV) or the MPRV feature in the mstatus register is used to execute a load or store as though V=1.

9.3. Hypervisor Instructions

The hypervisor extension adds virtual-machine load and store instructions and two privileged fence instructions.

9.3.1. Hypervisor Virtual-Machine Load and Store Instructions

The hypervisor virtual-machine load and store instructions are valid only in M-mode or HS-mode, or in U-mode when hstatus.HU=1. Each instruction performs an explicit memory access as though V=1; i.e., with the address translation and protection, and the endianness, that apply to memory accesses in either VS-mode or VU-mode. Field SPVP of hstatus controls the privilege level of the access. The explicit memory access is done as though in VU-mode when SPVP=0, and as though in VS-mode when SPVP=1. As usual when V=1, two-stage address translation is applied, and the HS-level sstatus.SUM is ignored. HS-level sstatus.MXR makes execute-only pages readable for both stages of address translation (VS-stage and G-stage), whereas vsstatus.MXR affects only the first translation stage (VS-stage).

For every RV32I or RV64I load instruction, LB, LBU, LH, LHU, LW, LWU, and LD, there is a corresponding virtual-machine load instruction: HLV.B, HLV.BU, HLV.H, HLV.HU, HLV.W, HLV.WU, and HLV.D. For every RV32I or RV64I store instruction, SB, SH, SW, and SD, there is a corresponding virtual-machine store instruction: HSV.B, HSV.H, HSV.W, and HSV.D. Instructions HLV.WU, HLV.D, and HSV.D are not valid for RV32, of course.

Instructions HLVX.HU and HLVX.WU are the same as HLV.HU and HLV.WU, except that execute permission takes the place of read permission during address translation. That is, the memory being read must be executable in both stages of address translation, but read permission is not required. For the supervisor physical address that results from address translation, the supervisor physical memory attributes must grant both execute and read permissions. (The supervisor physical memory attributes are the machine’s physical memory attributes as modified by physical memory protection, Section 3.7, for supervisor level.)

HLVX cannot override machine-level physical memory protection (PMP), so attempting to read memory that PMP designates as execute-only still results in an access-fault exception. Although HLVX instructions’ explicit memory accesses require execute permissions, they still raise the same exceptions as other load instructions, rather than raising fetch exceptions instead.

HLVX.WU is valid for RV32, even though LWU and HLV.WU are not. (For RV32, HLVX.WU can be considered a variant of HLV.W, as sign extension is irrelevant for 32-bit values.)

Attempts to execute a virtual-machine load/store instruction (HLV, HLVX, or HSV) when V=1 cause a virtual-instruction exception. Attempts to execute one of these same instructions from U-mode when hstatus.HU=0 cause an illegal-instruction exception.

9.3.2. Hypervisor Memory-Management Fence Instructions

The hypervisor memory-management fence instructions, HFENCE.VVMA and HFENCE.GVMA, perform a function similar to SFENCE.VMA (Section 5.2.1), except applying to the VS-level memory-management data structures controlled by CSR vsatp (HFENCE.VVMA) or the guest-physical memory-management data structures controlled by CSR hgatp (HFENCE.GVMA). Instruction SFENCE.VMA applies only to the memory-management data structures controlled by the current satp (either the HS-level satp when V=0 or vsatp when V=1).

HFENCE.VVMA is valid only in M-mode or HS-mode. Its effect is much the same as temporarily entering VS-mode and executing SFENCE.VMA. Executing an HFENCE.VVMA guarantees that any previous stores already visible to the current hart are ordered before all implicit reads by that hart done for VS-stage address translation for instructions that

• are subsequent to the HFENCE.VVMA, and

• execute when hgatp.VMID has the same setting as it did when HFENCE.VVMA executed.

Implicit reads need not be ordered when hgatp.VMID is different than at the time HFENCE.VVMA executed. If operand rs1≠x0, it specifies a single guest virtual address, and if operand rs2≠x0, it specifies a single guest address-space identifier (ASID).

An HFENCE.VVMA instruction applies only to a single virtual machine, identified by the setting of hgatp.VMID when HFENCE.VVMA executes.

When rs2≠x0, bits XLEN-1:ASIDMAX of the value held in rs2 are reserved for future standard use. Until their use is defined by a standard extension, they should be zeroed by software and ignored by current implementations. Furthermore, if ASIDLEN < ASIDMAX, the implementation shall ignore bits ASIDMAX-1:ASIDLEN of the value held in rs2.

Simpler implementations of HFENCE.VVMA can ignore the guest virtual address in rs1 and the guest ASID value in rs2, as well as hgatp.VMID, and always perform a global fence for the VS-level memory management of all virtual machines, or even a global fence for all memory-management data structures.

Neither mstatus.TVM nor hstatus.VTVM causes HFENCE.VVMA to trap.

HFENCE.GVMA is valid only in HS-mode when mstatus.TVM=0, or in M-mode (irrespective of mstatus.TVM). Executing an HFENCE.GVMA instruction guarantees that any previous stores already visible to the current hart are ordered before all implicit reads by that hart done for G-stage address translation for instructions that follow the HFENCE.GVMA. If operand rs1≠x0, it specifies a single guest physical address, shifted right by 2 bits, and if operand rs2≠x0, it specifies a single virtual machine identifier (VMID).

Conceptually, an implementation might contain two address-translation caches: one that maps guest virtual addresses to guest physical addresses, and another that maps guest physical addresses to supervisor physical addresses. HFENCE.GVMA need not flush the former cache, but it must flush entries from the latter cache that match the HFENCE.GVMA’s address and VMID arguments. More commonly, implementations contain address-translation caches that map guest virtual addresses directly to supervisor physical addresses, removing a level of indirection. For such implementations, any entry whose guest virtual address maps to a guest physical address that matches the HFENCE.GVMA’s address and VMID arguments must be flushed. Selectively flushing entries in this fashion requires tagging them with the guest physical address, which is costly, and so a common technique is to flush all entries that match the HFENCE.GVMA’s VMID argument, regardless of the address argument. Like for a guest physical address written to htval on a trap, a guest physical address specified in rs1 is shifted right by 2 bits to accommodate addresses wider than the current XLEN.

When rs2≠x0, bits XLEN-1:VMIDMAX of the value held in rs2 are reserved for future standard use. Until their use is defined by a standard extension, they should be zeroed by software and ignored by current implementations. Furthermore, if VMIDLEN < VMIDMAX, the implementation shall ignore bits VMIDMAX-1:VMIDLEN of the value held in rs2.

Simpler implementations of HFENCE.GVMA can ignore the guest physical address in rs1 and the VMID value in rs2 and always perform a global fence for the guest-physical memory management of all virtual machines, or even a global fence for all memory-management data structures.

If hgatp.MODE is changed for a given VMID, an HFENCE.GVMA with rs1=x0 (and rs2 set to either x0 or the VMID) must be executed to order subsequent guest translations with the MODE change—even if the old MODE or new MODE is Bare.

Attempts to execute HFENCE.VVMA or HFENCE.GVMA when V=1 cause a virtual-instruction exception, while attempts to do the same in U-mode cause an illegal-instruction exception. Attempting to execute HFENCE.GVMA in HS-mode when mstatus.TVM=1 also causes an illegal-instruction exception.

9.4. Machine-Level CSRs

The hypervisor extension augments or modifies machine CSRs mstatus, mstatush, mideleg, mip, and mie, and adds CSRs mtval2 and mtinst.

9.4.1. Machine Status Registers (mstatus and mstatush)

The hypervisor extension adds two fields, MPV and GVA, to the machine-level mstatus or mstatush CSR, and modifies the behavior of several existing mstatus fields. Figure 99 shows the modified mstatus register when the hypervisor extension is implemented and MXLEN=64. When MXLEN=32, the hypervisor extension adds MPV and GVA not to mstatus but to mstatush. Figure 100 shows the mstatush register when the hypervisor extension is implemented and MXLEN=32.

Figure 99. Machine status register (mstatus) fpr RV64 when the hypervisor extension is implemented.

Figure 100. Additional machine status register (mstatush) for RV32 when the hypervisor extension is implemented. The format of mstatus is unchanged for RV32.

The MPV bit (Machine Previous Virtualization Mode) is written by the implementation whenever a trap is taken into M-mode. Just as the MPP field is set to the (nominal) privilege mode at the time of the trap, the MPV bit is set to the value of the virtualization mode V at the time of the trap. When an MRET instruction is executed, the virtualization mode V is set to MPV, unless MPP=3, in which case V remains 0.

Field GVA (Guest Virtual Address) is written by the implementation whenever a trap is taken into M-mode. For any trap (breakpoint, address misaligned, access fault, page fault, or guest-page fault) that writes a guest virtual address to mtval, GVA is set to 1. For any other trap into M-mode, GVA is set to 0.

The TSR and TVM fields of mstatus affect execution only in HS-mode, not in VS-mode. The TW field affects execution in all modes except M-mode.

Setting TVM=1 prevents HS-mode from accessing hgatp or executing HFENCE.GVMA or HINVAL.GVMA, but has no effect on accesses to vsatp or instructions HFENCE.VVMA or HINVAL.VVMA.

TVM exists in mstatus to allow machine-level software to modify the address translations managed by a supervisor-level OS, usually for the purpose of inserting another stage of address translation below that controlled by the OS. The instruction traps enabled by TVM=1 permit machine level to co-opt both satp and hgatp and substitute shadow page tables that merge the OS’s chosen page translations with M-level’s lower-stage translations, all without the OS being aware. M-level software needs this ability not only to emulate the hypervisor extension if not already supported, but also to emulate any future RISC-V extensions that may modify or add address translation stages, perhaps, for example, to improve support for nested hypervisors, i.e., running hypervisors atop other hypervisors. However, setting TVM=1 does not cause traps for accesses to vsatp or instructions HFENCE.VVMA or HINVAL.VVMA, or for any actions taken in VS-mode, because M-level software is not expected to need to involve itself in VS-stage address translation. For virtual machines, it should be sufficient, and in all likelihood faster as well, to leave VS-stage address translation alone and merge all other translation stages into G-stage shadow page tables controlled by hgatp. This assumption does place some constraints on possible future RISC-V extensions that current machines will be able to emulate efficiently.

The hypervisor extension changes the behavior of the Modify Privilege field, MPRV, of mstatus. When MPRV=0, translation and protection behave as normal. When MPRV=1, explicit memory accesses are translated and protected, and endianness is applied, as though the current virtualization mode were set to MPV and the current nominal privilege mode were set to MPP. Table 31 enumerates the cases.

Table 31. Effect of MPRV on the translation and protection of explicit memory accesses.

MPRV	MPV	MPP	Effect
0	-	-	Normal access; current privilege mode applies.
1	0	0	U-level access with HS-level translation and protection only.
1	0	1	HS-level access with HS-level translation and protection only.
1	-	3	M-level access with no translation.
1	1	0	VU-level access with two-stage translation and protection. The HS-level MXR bit makes any executable page readable. vsstatus.MXR makes readable those pages marked executable at the VS translation stage, but only if readable at the guest-physical translation stage.
1	1	1	VS-level access with two-stage translation and protection. The HS-level MXR bit makes any executable page readable. vsstatus.MXR makes readable those pages marked executable at the VS translation stage, but only if readable at the guest-physical translation stage. vsstatus.SUM applies instead of the HS-level SUM bit.

Table 31. Effect of MPRV on the translation and protection of explicit memory accesses. {#h-mprv}

MPRV does not affect the virtual-machine load/store instructions, HLV, HLVX, and HSV. The explicit loads and stores of these instructions always act as though V=1 and the nominal privilege mode were hstatus.SPVP, overriding MPRV.

The mstatus register is a superset of the HS-level sstatus register but is not a superset of vsstatus.

9.4.2. Machine Interrupt Delegation Register (mideleg)

When the hypervisor extension is implemented, bits 10, 6, and 2 of mideleg (corresponding to the standard VS-level interrupts) are each read-only one. Furthermore, if any guest external interrupts are implemented (GEILEN is nonzero), bit 12 of mideleg (corresponding to supervisor-level guest external interrupts) is also read-only one. VS-level interrupts and guest external interrupts are always delegated past M-mode to HS-mode.

For bits of mideleg that are zero, the corresponding bits in hideleg, hip, and hie are read-only zeros.

9.4.3. Machine Interrupt Registers (mip and mie)

The hypervisor extension gives registers mip and mie additional active bits for the hypervisor-added interrupts. Figure 101 and Figure 102 show the standard portions (bits 15:0) of registers mip and mie when the hypervisor extension is implemented.

Figure 101. Standard portion (bits 15:0) of mip.

Figure 102. Standard portion (bits 15:0) of mie.

Bits SGEIP, VSEIP, VSTIP, and VSSIP in mip are aliases for the same bits in hypervisor CSR hip, while SGEIE, VSEIE, VSTIE, and VSSIE in mie are aliases for the same bits in hie.

9.4.4. Machine Second Trap Value Register (mtval2)

The mtval2 register is an MXLEN-bit read/write register formatted as shown in Figure 103. When a trap is taken into M-mode, mtval2 is written with additional exception-specific information, alongside mtval, to assist software in handling the trap.

Figure 103. Machine second trap value register (mtval2).

When a guest-page-fault trap is taken into M-mode, mtval2 is written with either zero or the guest physical address that faulted, shifted right by 2 bits. For other traps, mtval2 is set to zero, but a future standard or extension may redefine mtval2’s setting for other traps.

If a guest-page fault is due to an implicit memory access during first-stage (VS-stage) address translation, a guest physical address written to mtval2 is that of the implicit memory access that faulted. Additional information is provided in CSR mtinst to disambiguate such situations.

Otherwise, for misaligned loads and stores that cause guest-page faults, a nonzero guest physical address in mtval2 corresponds to the faulting portion of the access as indicated by the virtual address in mtval. For instruction guest-page faults on systems with variable-length instructions, a nonzero mtval2 corresponds to the faulting portion of the instruction as indicated by the virtual address in mtval.

mtval2 is a WARL register that must be able to hold zero and may be capable of holding only an arbitrary subset of other 2-bit-shifted guest physical addresses, if any.

9.4.5. Machine Trap Instruction Register (mtinst)

The mtinst register is an MXLEN-bit read/write register formatted as shown in Figure 104. When a trap is taken into M-mode, mtinst is written with a value that, if nonzero, provides information about the instruction that trapped, to assist software in handling the trap. The values that may be written to mtinst on a trap are documented in Section 9.6.3.

Figure 104. Machine trap instruction register (mtinst).

mtinst is a WARL register that need only be able to hold the values that the implementation may automatically write to it on a trap.

9.5. Two-Stage Address Translation

Whenever the current virtualization mode V is 1, two-stage address translation and protection is in effect. For any virtual memory access, the original virtual address is converted in the first stage by VS-level address translation, as controlled by the vsatp register, into a guest physical address. The guest physical address is then converted in the second stage by guest physical address translation, as controlled by the hgatp register, into a supervisor physical address. The two stages are known also as VS-stage and G-stage translation. Although there is no option to disable two-stage address translation when V=1, either stage of translation can be effectively disabled by zeroing the corresponding vsatp or hgatp register.

The vsstatus field MXR, which makes execute-only pages readable, only overrides VS-stage page protection. Setting MXR at VS-level does not override guest-physical page protections. Setting MXR at HS-level, however, overrides both VS-stage and G-stage execute-only permissions.

When V=1, memory accesses that would normally bypass address translation are subject to G-stage address translation alone. This includes memory accesses made in support of VS-stage address translation, such as reads and writes of VS-level page tables.

Machine-level physical memory protection applies to supervisor physical addresses and is in effect regardless of virtualization mode.

9.5.1. Guest Physical Address Translation

The mapping of guest physical addresses to supervisor physical addresses is controlled by CSR hgatp (Section 9.2.10).

When the address translation scheme selected by the MODE field of hgatp is Bare, guest physical addresses are equal to supervisor physical addresses without modification, and no memory protection applies in the trivial translation of guest physical addresses to supervisor physical addresses.

When hgatp.MODE specifies a translation scheme of Sv32x4, Sv39x4, Sv48x4, or Sv57x4, G-stage address translation is a variation on the usual page-based virtual address translation scheme of Sv32, Sv39, Sv48, or Sv57, respectively. In each case, the size of the incoming address is widened by 2 bits (to 34, 41, 50, or 59 bits). To accommodate the 2 extra bits, the root page table (only) is expanded by a factor of four to be 16 KiB instead of the usual 4 KiB. Matching its larger size, the root page table also must be aligned to a 16 KiB boundary instead of the usual 4 KiB page boundary. Except as noted, all other aspects of Sv32, Sv39, Sv48, or Sv57 are adopted unchanged for G-stage translation. Non-root page tables and all page table entries (PTEs) have the same formats as documented in Section 5.3, Section 5.4, Section 5.5, and Section 5.6.

For Sv32x4, an incoming guest physical address is partitioned into a virtual page number (VPN) and page offset as shown in Figure 105. This partitioning is identical to that for an Sv32 virtual address as depicted in Figure 55, except with 2 more bits at the high end in VPN[1]. (Note that the fields of a partitioned guest physical address also correspond one-for-one with the structure that Sv32 assigns to a physical address, depicted in Figure 55.)

Figure 105. Sv32x4 virtual address (guest physical address).

For Sv39x4, an incoming guest physical address is partitioned as shown in Figure 106. This partitioning is identical to that for an Sv39 virtual address as depicted in Figure 58, except with 2 more bits at the high end in VPN[2]. Address bits 63:41 must all be zeros, or else a guest-page-fault exception occurs.

Figure 106. Sv39x4 virtual address (guest physical address).

For Sv48x4, an incoming guest physical address is partitioned as shown in Figure 107. This partitioning is identical to that for an Sv48 virtual address as depicted in Figure 61, except with 2 more bits at the high end in VPN[3]. Address bits 63:50 must all be zeros, or else a guest-page-fault exception occurs.

Figure 107. Sv48x4 virtual address (guest physical address).

For Sv57x4, an incoming guest physical address is partitioned as shown in Figure 108. This partitioning is identical to that for an Sv57 virtual address as depicted in Figure 64, except with 2 more bits at the high end in VPN[4]. Address bits 63:59 must all be zeros, or else a guest-page-fault exception occurs.

Figure 108. Sv57x4 virtual address (guest physical address).

The page-based G-stage address translation scheme for RV32, Sv32x4, is defined to support a 34-bit guest physical address so that an RV32 hypervisor need not be limited in its ability to virtualize real 32-bit RISC-V machines, even those with 33-bit or 34-bit physical addresses. This may include the possibility of a machine virtualizing itself, if it happens to use 33-bit or 34-bit physical addresses. Multiplying the size and alignment of the root page table by a factor of four is the cheapest way to extend Sv32 to cover a 34-bit address. The possible wastage of 12 KiB for an unnecessarily large root page table is expected to be of negligible consequence for most (maybe all) real uses. A consistent ability to virtualize machines having as much as four times the physical address space as virtual address space is believed to be of some utility also for RV64. For a machine implementing 39-bit virtual addresses (Sv39), for example, this allows the hypervisor extension to support up to a 41-bit guest physical address space without either necessitating hardware support for 48-bit virtual addresses (Sv48) or falling back to emulating the larger address space using shadow page tables.

The conversion of an Sv32x4, Sv39x4, Sv48x4, or Sv57x4 guest physical address is accomplished with the same algorithm used for Sv32, Sv39, Sv48, or Sv57, as presented in Section 5.3.2, except that:

• hgatp substitutes for the usual satp;

• for the translation to begin, the effective privilege mode must be VS-mode or VU-mode;

• when checking the U bit, the current privilege mode is always taken to be U-mode; and

• guest-page-fault exceptions are raised instead of regular page-fault exceptions.

For G-stage address translation, all memory accesses (including those made to access data structures for VS-stage address translation) are considered to be user-level accesses, as though executed in U-mode. Access type permissions—readable, writable, or executable—are checked during G-stage translation the same as for VS-stage translation. For a memory access made to support VS-stage address translation (such as to read/write a VS-level page table), permissions are checked as though for a load or store, not for the original access type. However, any exception is always reported for the original access type (instruction, load, or store/AMO).

The G bit in all G-stage PTEs is reserved for future standard use. Until its use is defined by a standard extension, it should be cleared by software for forward compatibility, and must be ignored by hardware.

G-stage address translation uses the identical format for PTEs as regular address translation, even including the U bit, due to the possibility of sharing some (or all) page tables between G-stage translation and regular HS-level address translation. Regardless of whether this usage will ever become common, we chose not to preclude it.

9.5.2. Guest-Page Faults

Guest-page-fault traps may be delegated from M-mode to HS-mode under the control of CSR medeleg, but cannot be delegated to other privilege modes. On a guest-page fault, CSR mtval or stval is written with the faulting guest virtual address as usual, and mtval2 or htval is written either with zero or with the faulting guest physical address, shifted right by 2 bits. CSR mtinst or htinst may also be written with information about the faulting instruction or other reason for the access, as explained in Section 9.6.3.

When an instruction fetch or a misaligned memory access straddles a page boundary, two different address translations are involved. When a guest-page fault occurs in such a circumstance, the faulting virtual address written to mtval/stval is the same as would be required for a regular page fault. Thus, the faulting virtual address may be a page-boundary address that is higher than the instruction’s original virtual address, if the byte at that page boundary is among the accessed bytes.

When a guest-page fault is not due to an implicit memory access for VS-stage address translation, a nonzero guest physical address written to mtval2/htval shall correspond to the exact virtual address written to mtval/stval.

9.5.3. Memory-Management Fences

The behavior of the SFENCE.VMA instruction is affected by the current virtualization mode V. When V=0, the virtual-address argument is an HS-level virtual address, and the ASID argument is an HS-level ASID. The instruction orders stores only to HS-level address-translation structures with subsequent HS-level address translations.

When V=1, the virtual-address argument to SFENCE.VMA is a guest virtual address within the current virtual machine, and the ASID argument is a VS-level ASID within the current virtual machine. The current virtual machine is identified by the VMID field of CSR hgatp, and the effective ASID can be considered to be the combination of this VMID with the VS-level ASID. The SFENCE.VMA instruction orders stores only to the VS-level address-translation structures with subsequent VS-stage address translations for the same virtual machine, i.e., only when hgatp.VMID is the same as when the SFENCE.VMA executed.

Hypervisor instructions HFENCE.VVMA and HFENCE.GVMA provide additional memory-management fences to complement SFENCE.VMA. These instructions are described in Section 9.3.2.

Section 3.7.2 discusses the intersection between physical memory protection (PMP) and page-based address translation. It is noted there that, when PMP settings are modified in a manner that affects either the physical memory that holds page tables or the physical memory to which page tables point, M-mode software must synchronize the PMP settings with the virtual memory system. For HS-level address translation, this is accomplished by executing in M-mode an SFENCE.VMA instruction with rs1=x0 and rs2=x0, after the PMP CSRs are written. Synchronization with G-stage and VS-stage data structures is also needed. Executing an HFENCE.GVMA instruction with rs1=x0 and rs2=x0 suffices to flush all G-stage or VS-stage address-translation cache entries that have cached PMP settings corresponding to the final translated supervisor physical address. An HFENCE.VVMA instruction is not required.

9.6. Traps

9.6.1. Trap Cause Codes

The hypervisor extension augments the trap cause encoding. Table 32 lists the possible M-mode and HS-mode trap cause codes when the hypervisor extension is implemented. Codes are added for VS-level interrupts (interrupts 2, 6, 10), for supervisor-level guest external interrupts (interrupt 12), for virtual-instruction exceptions (exception 22), and for guest-page faults (exceptions 20, 21, 23). Furthermore, environment calls from VS-mode are assigned cause 10, whereas those from HS-mode or S-mode use cause 9 as usual.

Table 32. Machine and supervisor cause register (mcause and scause) values when the hypervisor extension is implemented.

Interrupt	Exception Code	Description
1 1 1 1	0 1 2 3	Reserved Supervisor software interrupt Virtual supervisor software interrupt Machine software interrupt
1 1 1 1	4 5 6 7	Reserved Supervisor timer interrupt Virtual supervisor timer interrupt Machine timer interrupt
1 1 1 1	8 9 10 11	Reserved Supervisor external interrupt Virtual supervisor external interrupt Machine external interrupt
1 1 1	12 13-15 ≥16	Supervisor guest external interrupt Reserved Designated for platform or custom use
0 0 0 0 0 0 0 0 0 0 0 0 0 0 0 0 0 0 0 0 0 0 0 0 0	0 1 2 3 4 5 6 7 8 9 10 11 12 13 14 15 16-19 20 21 22 23 24-31 32-47 48-63 ≥64	Instruction address misaligned Instruction access fault Illegal instruction Breakpoint Load address misaligned Load access fault Store/AMO address misaligned Store/AMO access fault Environment call from U-mode or VU-mode Environment call from HS-mode Environment call from VS-mode Environment call from M-mode Instruction page fault Load page fault Reserved Store/AMO page fault Reserved Instruction guest-page fault Load guest-page fault Virtual instruction Store/AMO guest-page fault Designated for custom use Reserved Designated for custom use Reserved

Table 32. Machine and supervisor cause register (mcause and scause) values when the hypervisor extension is implemented. {#hcauses}

HS-mode and VS-mode ECALLs use different cause values so they can be delegated separately.

When V=1, a virtual-instruction exception (code 22) is normally raised instead of an illegal-instruction exception if the attempted instruction is HS-qualified but is prevented from executing when V=1 either due to insufficient privilege or because the instruction is expressly disabled by a supervisor or hypervisor CSR such as scounteren or hcounteren. An instruction is HS-qualified if it would be valid to execute in HS-mode (for some values of the instruction’s register operands), assuming fields TSR and TVM of CSR mstatus are both zero.

A special rule applies for CSR instructions that access 32-bit high-half CSRs such as cycleh and htimedeltah. When V=1 and XLEN=32, an invalid attempt to access a high-half CSR raises a virtual-instruction exception instead of an illegal-instruction exception if the same CSR instruction for the corresponding low-half CSR (e.g.cycle or htimedelta) is HS-qualified.

When XLEN>32, an attempt to access a high-half CSR always raises an illegal-instruction exception.

Specifically, a virtual-instruction exception is raised for the following cases:

• in VS-mode, attempts to access a non-high-half counter CSR when the corresponding bit in hcounteren is 0 and the same bit in mcounteren is 1;

• in VS-mode, if XLEN=32, attempts to access a high-half counter CSR when the corresponding bit in hcounteren is 0 and the same bit in mcounteren is 1;

• in VU-mode, attempts to access a non-high-half counter CSR when the corresponding bit in either hcounteren or scounteren is 0 and the same bit in mcounteren is 1;

• in VU-mode, if XLEN=32, attempts to access a high-half counter CSR when the corresponding bit in either hcounteren or scounteren is 0 and the same bit in mcounteren is 1;

• in VS-mode or VU-mode, attempts to execute a hypervisor instruction (HLV, HLVX, HSV, or HFENCE);

• in VS-mode or VU-mode, attempts to access an implemented non-high-half hypervisor CSR or VS CSR when the same access (read/write) would be allowed in HS-mode, assuming mstatus.TVM=0;

• in VS-mode or VU-mode, if XLEN=32, attempts to access an implemented high-half hypervisor CSR or high-half VS CSR when the same access (read/write) to the CSR"s low-half partner would be allowed in HS-mode, assuming mstatus.TVM=0;

• in VU-mode, attempts to execute WFI when mstatus.TW=0, or to execute a supervisor instruction (SRET or SFENCE);

• in VU-mode, attempts to access an implemented non-high-half supervisor CSR when the same access (read/write) would be allowed in HS-mode, assuming mstatus.TVM=0;

• in VU-mode, if XLEN=32, attempts to access an implemented high-half supervisor CSR when the same access to the CSR’s low-half partner would be allowed in HS-mode, assuming mstatus.TVM=0;

• in VS-mode, attempts to execute WFI when hstatus.VTW=1 and mstatus.TW=0, unless the instruction completes within an implementation-specific, bounded time;

• in VS-mode, attempts to execute SRET when hstatus.VTSR=1; and

• in VS-mode, attempts to execute an SFENCE.VMA or SINVAL.VMA instruction or to access satp, when hstatus.VTVM=1.

Other extensions to the RISC-V Privileged Architecture may add to the set of circumstances that cause a virtual-instruction exception when V=1.

On a virtual-instruction trap, mtval or stval is written the same as for an illegal-instruction trap.

It is not unusual that hypervisors must emulate the instructions that raise virtual-instruction exceptions, to support nested hypervisors or for other reasons. Machine level is expected ordinarily to delegate virtual-instruction traps directly to HS-level, whereas illegal-instruction traps are likely to be processed first in M-mode before being conditionally delegated (by software) to HS-level. Consequently, virtual-instruction traps are expected typically to be handled faster than illegal-instruction traps. When not emulating the trapping instruction, a hypervisor should convert a virtual-instruction trap into an illegal-instruction exception for the guest virtual machine. Because TSR and TVM in mstatus are intended to impact only S-mode (HS-mode), they are ignored for determining exceptions in VS-mode.

Table 33. Synchronous exception priority when the hypervisor extension is implemented.

Priority	Exc.Code	Description
Highest	3	Instruction address breakpoint
	12, 20, 1	During instruction address translation:    First encountered page fault, guest-page fault, or access fault
	1	With physical address for instruction:    Instruction access fault
	2 22 0 8, 9, 10, 11 3 3	Illegal instruction Virtual instruction Instruction address misaligned Environment call Environment break    Load/store/AMO address breakpoint
	4,6	Optionally:    Load/store/AMO address misaligned
	13, 15, 21, 23, 5, 7	During address translation for an explicit memory access:    First encountered page fault, guest-page fault, or access fault
	5, 7	With physical address for an explicit memory access:    Load/store/AMO access fault
Lowest	4, 6	If not higher priority:    Load/store/AMO address misaligned

Table 33. Synchronous exception priority when the hypervisor extension is implemented. {#HSyncExcPrio}

If an instruction may raise multiple synchronous exceptions, the decreasing priority order of Table 33 indicates which exception is taken and reported in mcause or scause.

9.6.2. Trap Entry

When a trap occurs in HS-mode or U-mode, it goes to M-mode, unless delegated by medeleg or mideleg, in which case it goes to HS-mode. When a trap occurs in VS-mode or VU-mode, it goes to M-mode, unless delegated by medeleg or mideleg, in which case it goes to HS-mode, unless further delegated by hedeleg or hideleg, in which case it goes to VS-mode.

When a trap is taken into M-mode, virtualization mode V gets set to 0, and fields MPV and MPP in mstatus (or mstatush) are set according to Table 34. A trap into M-mode also writes fields GVA, MPIE, and MIE in mstatus/mstatush and writes CSRs mepc, mcause, mtval, mtval2, and mtinst.

Table 34. Value of mstatus/mstatush fields MPV and MPP after a trap into M-mode. Upon trap return, MPV is ignored when MPP=3.

Previous Mode	MPV	MPP
U-mode HS-mode M-mode	0 0 0	0 1 3
VU-mode VS-mode	1 1	0 1

Table 34. Value of mstatus/mstatush fields MPV and MPP after a trap into M-mode. Upon trap return, MPV is ignored when MPP=3. {#h-mpp}

When a trap is taken into HS-mode, virtualization mode V is set to 0, and hstatus.SPV and sstatus.SPP are set according to Table 35. If V was 1 before the trap, field SPVP in hstatus is set the same as sstatus.SPP; otherwise, SPVP is left unchanged. A trap into HS-mode also writes field GVA in hstatus, fields SPIE and SIE in sstatus, and CSRs sepc, scause, stval, htval, and htinst.

Table 35. Value of hstatus field SPV and sstatus field SPP after a trap into HS-mode.

Previous Mode	SPV	SPP
U-mode HS-mode	0 0	0 1
VU-mode VS-mode	1 1	0 1

Table 35. Value of hstatus field SPV and sstatus field SPP after a trap into HS-mode. {#h-spp}

When a trap is taken into VS-mode, vsstatus.SPP is set according to Table 36. Register hstatus and the HS-level sstatus are not modified, and the virtualization mode V remains 1. A trap into VS-mode also writes fields SPIE and SIE in vsstatus and writes CSRs vsepc, vscause, and vstval.

Table 36. Value of vsstatus field SPP after a trap into VS-mode.

Previous Mode	SPP
VU-mode VS-mode	0 1

Table 36. Value of vsstatus field SPP after a trap into VS-mode. {#h-vspp}

9.6.3. Transformed Instruction or Pseudoinstruction for mtinst or htinst

On any trap into M-mode or HS-mode, one of these values is written automatically into the appropriate trap instruction CSR, mtinst or htinst:

• zero;

• a transformation of the trapping instruction;

• a custom value (allowed only if the trapping instruction is non-standard); or

• a special pseudoinstruction.

Except when a pseudoinstruction value is required (described later), the value written to mtinst or htinst may always be zero, indicating that the hardware is providing no information in the register for this particular trap.

The value written to the trap instruction CSR serves two purposes. The first is to improve the speed of instruction emulation in a trap handler, partly by allowing the handler to skip loading the trapping instruction from memory, and partly by obviating some of the work of decoding and executing the instruction. The second purpose is to supply, via pseudoinstructions, additional information about guest-page-fault exceptions caused by implicit memory accesses done for VS-stage address translation. A transformation of the trapping instruction is written instead of simply a copy of the original instruction in order to minimize the burden for hardware yet still provide to a trap handler the information needed to emulate the instruction. An implementation may at any time reduce its effort by substituting zero in place of the transformed instruction.

On an interrupt, the value written to the trap instruction register is always zero. On a synchronous exception, if a nonzero value is written, one of the following shall be true about the value:

• Bit 0 is 1, and replacing bit 1 with 1 makes the value into a valid encoding of a standard instruction.

  In this case, the instruction that trapped is the same kind as indicated by the register value, and the register value is the transformation of the trapping instruction, as defined later. For example, if bits 1:0 are binary 11 and the register value is the encoding of a standard LW (load word) instruction, then the trapping instruction is LW, and the register value is the transformation of the trapping LW instruction.

• Bit 0 is 1, and replacing bit 1 with 1 makes the value into an instruction encoding that is explicitly designated for a custom instruction (not an unused reserved encoding).

  This is a custom value. The instruction that trapped is a non-standard instruction. The interpretation of a custom value is not otherwise specified by this standard.

• The value is one of the special pseudoinstructions defined later, all of which have bits 1:0 equal to 00.

These three cases exclude a large number of other possible values, such as all those having bits 1:0 equal to binary 10. A future standard or extension may define additional cases, thus allowing values that are currently excluded. Software may safely treat an unrecognized value in a trap instruction register the same as zero.

To be forward-compatible with future revisions of this standard, software that interprets a nonzero value from mtinst or htinst must fully verify that the value conforms to one of the cases listed above. For instance, for RV64, discovering that bits 6:0 of mtinst are 0000011 and bits 14:12 are 010 is not sufficient to establish that the first case applies and the trapping instruction is a standard LW instruction; rather, software must also confirm that bits 63:32 of mtinst are all zeros. A future standard might define new values for 64-bit mtinst that are nonzero in bits 63:32 yet may coincidentally have in bits 31:0 the same bit patterns as standard RV64 instructions. Unlike for standard instructions, there is no requirement that the instruction encoding of a custom value be of the same ``kind'' as the instruction that trapped (or even have any correlation with the trapping instruction).

Table 37 shows the values that may be automatically written to the trap instruction register for each standard exception cause. For exceptions that prevent the fetching of an instruction, only zero or a pseudoinstruction value may be written. A custom value may be automatically written only if the instruction that traps is non-standard. A future standard or extension may permit other values to be written, chosen from the set of allowed values established earlier.

Table 37. Values that may be automatically written to the trap instruction register (mtinst or htinst) on an exception trap.

Exception	Zero	Transformed Standard Instruction	Custom Value	Pseudo-instruction Value
Instruction address misaligned	Yes	No	Yes	No
Instruction access fault Illegal instruction Breakpoint Virtual instruction	Yes Yes Yes Yes	No No No No	No No Yes Yes	No No No No
Load address misaligned Load access fault Store/AMO address misaligned Store/AMO access fault	Yes Yes Yes Yes	Yes Yes Yes Yes	Yes Yes Yes Yes	No No No No
Environment call	Yes	No	Yes	No
Instruction page fault Load page fault Store/AMO page fault	Yes Yes Yes	No Yes Yes	No Yes Yes	No No No
Instruction guest-page fault Load guest-page fault Store/AMO guest-page fault	Yes Yes Yes	No Yes Yes	No Yes Yes	Yes Yes Yes

Table 37. Values that may be automatically written to the trap instruction register (mtinst or htinst) on an exception trap. {#tinst-values}

As enumerated in the table, a synchronous exception may write to the trap instruction register a standard transformation of the trapping instruction only for exceptions that arise from explicit memory accesses (from loads, stores, and AMO instructions). Accordingly, standard transformations are currently defined only for these memory-access instructions. If a synchronous trap occurs for a standard instruction for which no transformation has been defined, the trap instruction register shall be written with zero (or, under certain circumstances, with a special pseudoinstruction value).

For a standard load instruction that is not a compressed instruction and is one of LB, LBU, LH, LHU, LW, LWU, LD, FLW, FLD, FLQ, or FLH, the transformed instruction has the format shown in Figure 109.

Figure 109. Transformed noncompressed load instruction (LB, LBU, LH, LHU, LW, LWU, LD, FLW, FLD, FLQ, or FLH). Fields funct3, rd, and opcode are the same as the trapping load instruction.

For a standard store instruction that is not a compressed instruction and is one of SB, SH, SW, SD, FSW, FSD, FSQ, or FSH, the transformed instruction has the format shown in Figure 110.

Figure 110. Transformed noncompressed store instruction (SB, SH, SW, SD, FSW, FSD, FSQ, or FSH). Fields rs2, funct3, and opcode are the same as the trapping store instruction.

For a standard atomic instruction (load-reserved, store-conditional, or AMO instruction), the transformed instruction has the format shown in Figure 111.

Figure 111. Transformed atomic instruction (load-reserved, store-conditional, or AMO instruc-tion). All fields are the same as the trapping instruction except bits 19:15, Addr. Offset.

For a standard virtual-machine load/store instruction (HLV, HLVX, or HSV), the transformed instruction has the format shown in Figure 112.

Figure 112. Transformed virtual-machine load/store instruction (HLV, HLVX, HSV). All fields are the same as the trapping instruction except bits 19:15, Addr. Offset

In all the transformed instructions above, the Addr. Offset field that replaces the instruction’s rs1 field in bits 19:15 is the positive difference between the faulting virtual address (written to mtval or stval) and the original virtual address. This difference can be nonzero only for a misaligned memory access. Note also that, for basic loads and stores, the transformations replace the instruction’s immediate offset fields with zero.

For a standard compressed instruction (16-bit size), the transformed instruction is found as follows:

1. Expand the compressed instruction to its 32-bit equivalent.

2. Transform the 32-bit equivalent instruction.

3. Replace bit 1 with a 0.

Bits 1:0 of a transformed standard instruction will be binary 01 if the trapping instruction is compressed and 11 if not.

In decoding the contents of mtinst or htinst, once software has determined that the register contains the encoding of a standard basic load (LB, LBU, LH, LHU, LW, LWU, LD, FLW, FLD, FLQ, or FLH) or basic store (SB, SH, SW, SD, FSW, FSD, FSQ, or FSH), it is not necessary to confirm also that the immediate offset fields (31:25, and 24:20 or 11:7) are zeros. The knowledge that the register’s value is the encoding of a basic load/store is sufficient to prove that the trapping instruction is of the same kind. A future version of this standard may add information to the fields that are currently zeros. However, for backwards compatibility, any such information will be for performance purposes only and can safely be ignored.

For guest-page faults, the trap instruction register is written with a special pseudoinstruction value if: (a) the fault is caused by an implicit memory access for VS-stage address translation, and (b) a nonzero value (the faulting guest physical address) is written to mtval2 or htval. If both conditions are met, the value written to mtinst or htinst must be taken from Table 38; zero is not allowed.

Table 38. Special pseudoinstruction values for guest-page faults. The RV32 values are used when VSXLEN=32, and the TV64 values when VSXLEN=64.

Value	Meaning
0x00002000 0x00002020	32-bit read for VS-stage address translation (RV32) 32-bit write for VS-stage address translation (RV32)
0x00003000 0x00003020	64-bit read for VS-stage address translation (RV64) 64-bit write for VS-stage address translation (RV64)

Table 38. Special pseudoinstruction values for guest-page faults. The RV32 values are used when VSXLEN=32, and the TV64 values when VSXLEN=64. {#pseudoinsts}

The defined pseudoinstruction values are designed to correspond closely with the encodings of basic loads and stores, as illustrated by Table 39.

Table 39. Standard instructions corresponding to the special psudoinstructions of Table 38.

Encoding	Instruction
0x00002003 0x00002023	lw x0,0(x0) sw x0,0(x0)
0x00003003 0x00003023	ld x0,0(x0) sd x0,0(x0)

Table 39. Standard instructions corresponding to the special psudoinstructions of Table 38. {#pseudoinsts-basis}

A write pseudoinstruction (0x00002020 or 0x00003020) is used for the case that the machine is attempting automatically to update bits A and/or D in VS-level page tables. All other implicit memory accesses for VS-stage address translation will be reads. If a machine never automatically updates bits A or D in VS-level page tables (leaving this to software), the write case will never arise. The fact that such a page table update must actually be atomic, not just a simple write, is ignored for the pseudoinstruction.

If the conditions that necessitate a pseudoinstruction value can ever occur for M-mode, then mtinst cannot be entirely read-only zero; and likewise for HS-mode and htinst. However, in that case, the trap instruction registers may minimally support only values 0 and 0x00002000 or 0x00003000, and possibly 0x00002020 or 0x00003020, requiring as few as one or two flip-flops in hardware, per register. There is no harm here in ignoring the atomicity requirement for page table updates, because a hypervisor is not expected in these circumstances to emulate an implicit memory access that fails. Rather, the hypervisor is given enough information about the faulting access to be able to make the memory accessible (e.g. by restoring a missing page of virtual memory) before resuming execution by retrying the faulting instruction.

9.6.4. Trap Return

The MRET instruction is used to return from a trap taken into M-mode. MRET first determines what the new privilege mode will be according to the values of MPP and MPV in mstatus or mstatush, as encoded in Table 34. MRET then in mstatus/mstatush sets MPV=0, MPP=0, MIE=MPIE, and MPIE=1. Lastly, MRET sets the privilege mode as previously determined, and sets pc=mepc.

The SRET instruction is used to return from a trap taken into HS-mode or VS-mode. Its behavior depends on the current virtualization mode.

When executed in M-mode or HS-mode (i.e., V=0), SRET first determines what the new privilege mode will be according to the values in hstatus.SPV and sstatus.SPP, as encoded in Table 35. SRET then sets hstatus.SPV=0, and in sstatus sets SPP=0, SIE=SPIE, and SPIE=1. Lastly, SRET sets the privilege mode as previously determined, and sets pc=sepc.

When executed in VS-mode (i.e., V=1), SRET sets the privilege mode according to Table 36, in vsstatus sets SPP=0, SIE=SPIE, and SPIE=1, and lastly sets pc=vsepc.

10. RISC-V Privileged Instruction Set Listings

This chapter presents instruction-set listings for all instructions defined in the RISC-V Privileged Architecture.

The instruction-set listings for unprivileged instructions, including the ECALL and EBREAK instructions, are provided in Volume I of this manual.

Figure 113. RISC-V Privileged Instructions

11. History

11.1. Research Funding at UC Berkeley

Development of the RISC-V architecture and implementations has been partially funded by the following sponsors.

• Par Lab: Research supported by Microsoft (Award #024263) and Intel (Award #024894) funding and by matching funding by U.C. Discovery (Award #DIG07-10227). Additional support came from Par Lab affiliates Nokia, NVIDIA, Oracle, and Samsung.

• Project Isis: DoE Award DE-SC0003624.

• ASPIRE Lab: DARPA PERFECT program, Award HR0011-12-2-0016. DARPA POEM program Award HR0011-11-C-0100. The Center for Future Architectures Research (C-FAR), a STARnet center funded by the Semiconductor Research Corporation. Additional support from ASPIRE industrial sponsor, Intel, and ASPIRE affiliates, Google, Huawei, Nokia, NVIDIA, Oracle, and Samsung.

The content of this paper does not necessarily reflect the position or the policy of the US government and no official endorsement should be inferred.

Bibliography

Goldberg, R. P. (1974). Survey of virtual machine research. Computer, 7(6), 34–45.

Navarro, J., Iyer, S., Druschel, P., & Cox, A. (2002). Practical, Transparent Operating System Support for Superpages. SIGOPS Oper. Syst. Rev., 36(SI), 89–104. doi.org/10.1145/844128.844138

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