xemu/accel/tcg/cputlb.c

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/*
* Common CPU TLB handling
*
* Copyright (c) 2003 Fabrice Bellard
*
* This library is free software; you can redistribute it and/or
* modify it under the terms of the GNU Lesser General Public
* License as published by the Free Software Foundation; either
* version 2 of the License, or (at your option) any later version.
*
* This library is distributed in the hope that it will be useful,
* but WITHOUT ANY WARRANTY; without even the implied warranty of
* MERCHANTABILITY or FITNESS FOR A PARTICULAR PURPOSE. See the GNU
* Lesser General Public License for more details.
*
* You should have received a copy of the GNU Lesser General Public
* License along with this library; if not, see <http://www.gnu.org/licenses/>.
*/
#include "qemu/osdep.h"
tcg: drop global lock during TCG code execution This finally allows TCG to benefit from the iothread introduction: Drop the global mutex while running pure TCG CPU code. Reacquire the lock when entering MMIO or PIO emulation, or when leaving the TCG loop. We have to revert a few optimization for the current TCG threading model, namely kicking the TCG thread in qemu_mutex_lock_iothread and not kicking it in qemu_cpu_kick. We also need to disable RAM block reordering until we have a more efficient locking mechanism at hand. Still, a Linux x86 UP guest and my Musicpal ARM model boot fine here. These numbers demonstrate where we gain something: 20338 jan 20 0 331m 75m 6904 R 99 0.9 0:50.95 qemu-system-arm 20337 jan 20 0 331m 75m 6904 S 20 0.9 0:26.50 qemu-system-arm The guest CPU was fully loaded, but the iothread could still run mostly independent on a second core. Without the patch we don't get beyond 32206 jan 20 0 330m 73m 7036 R 82 0.9 1:06.00 qemu-system-arm 32204 jan 20 0 330m 73m 7036 S 21 0.9 0:17.03 qemu-system-arm We don't benefit significantly, though, when the guest is not fully loading a host CPU. Signed-off-by: Jan Kiszka <jan.kiszka@siemens.com> Message-Id: <1439220437-23957-10-git-send-email-fred.konrad@greensocs.com> [FK: Rebase, fix qemu_devices_reset deadlock, rm address_space_* mutex] Signed-off-by: KONRAD Frederic <fred.konrad@greensocs.com> [EGC: fixed iothread lock for cpu-exec IRQ handling] Signed-off-by: Emilio G. Cota <cota@braap.org> [AJB: -smp single-threaded fix, clean commit msg, BQL fixes] Signed-off-by: Alex Bennée <alex.bennee@linaro.org> Reviewed-by: Richard Henderson <rth@twiddle.net> Reviewed-by: Pranith Kumar <bobby.prani@gmail.com> [PM: target-arm changes] Acked-by: Peter Maydell <peter.maydell@linaro.org>
2017-02-23 18:29:11 +00:00
#include "qemu/main-loop.h"
#include "cpu.h"
#include "exec/exec-all.h"
#include "exec/memory.h"
#include "exec/address-spaces.h"
#include "exec/cpu_ldst.h"
#include "exec/cputlb.h"
#include "exec/memory-internal.h"
#include "exec/ram_addr.h"
#include "tcg/tcg.h"
#include "qemu/error-report.h"
#include "exec/log.h"
#include "exec/helper-proto.h"
#include "qemu/atomic.h"
/* DEBUG defines, enable DEBUG_TLB_LOG to log to the CPU_LOG_MMU target */
/* #define DEBUG_TLB */
/* #define DEBUG_TLB_LOG */
#ifdef DEBUG_TLB
# define DEBUG_TLB_GATE 1
# ifdef DEBUG_TLB_LOG
# define DEBUG_TLB_LOG_GATE 1
# else
# define DEBUG_TLB_LOG_GATE 0
# endif
#else
# define DEBUG_TLB_GATE 0
# define DEBUG_TLB_LOG_GATE 0
#endif
#define tlb_debug(fmt, ...) do { \
if (DEBUG_TLB_LOG_GATE) { \
qemu_log_mask(CPU_LOG_MMU, "%s: " fmt, __func__, \
## __VA_ARGS__); \
} else if (DEBUG_TLB_GATE) { \
fprintf(stderr, "%s: " fmt, __func__, ## __VA_ARGS__); \
} \
} while (0)
#define assert_cpu_is_self(this_cpu) do { \
if (DEBUG_TLB_GATE) { \
g_assert(!cpu->created || qemu_cpu_is_self(cpu)); \
} \
} while (0)
/* run_on_cpu_data.target_ptr should always be big enough for a
* target_ulong even on 32 bit builds */
QEMU_BUILD_BUG_ON(sizeof(target_ulong) > sizeof(run_on_cpu_data));
/* We currently can't handle more than 16 bits in the MMUIDX bitmask.
*/
QEMU_BUILD_BUG_ON(NB_MMU_MODES > 16);
#define ALL_MMUIDX_BITS ((1 << NB_MMU_MODES) - 1)
/* flush_all_helper: run fn across all cpus
*
* If the wait flag is set then the src cpu's helper will be queued as
* "safe" work and the loop exited creating a synchronisation point
* where all queued work will be finished before execution starts
* again.
*/
static void flush_all_helper(CPUState *src, run_on_cpu_func fn,
run_on_cpu_data d)
{
CPUState *cpu;
CPU_FOREACH(cpu) {
if (cpu != src) {
async_run_on_cpu(cpu, fn, d);
}
}
}
/* statistics */
int tlb_flush_count;
/* This is OK because CPU architectures generally permit an
* implementation to drop entries from the TLB at any time, so
* flushing more entries than required is only an efficiency issue,
* not a correctness issue.
*/
static void tlb_flush_nocheck(CPUState *cpu)
{
CPUArchState *env = cpu->env_ptr;
/* The QOM tests will trigger tlb_flushes without setting up TCG
* so we bug out here in that case.
*/
if (!tcg_enabled()) {
return;
}
assert_cpu_is_self(cpu);
tlb_debug("(count: %d)\n", tlb_flush_count++);
tb_lock();
memset(env->tlb_table, -1, sizeof(env->tlb_table));
implementing victim TLB for QEMU system emulated TLB QEMU system mode page table walks are expensive. Taken by running QEMU qemu-system-x86_64 system mode on Intel PIN , a TLB miss and walking a 4-level page tables in guest Linux OS takes ~450 X86 instructions on average. QEMU system mode TLB is implemented using a directly-mapped hashtable. This structure suffers from conflict misses. Increasing the associativity of the TLB may not be the solution to conflict misses as all the ways may have to be walked in serial. A victim TLB is a TLB used to hold translations evicted from the primary TLB upon replacement. The victim TLB lies between the main TLB and its refill path. Victim TLB is of greater associativity (fully associative in this patch). It takes longer to lookup the victim TLB, but its likely better than a full page table walk. The memory translation path is changed as follows : Before Victim TLB: 1. Inline TLB lookup 2. Exit code cache on TLB miss. 3. Check for unaligned, IO accesses 4. TLB refill. 5. Do the memory access. 6. Return to code cache. After Victim TLB: 1. Inline TLB lookup 2. Exit code cache on TLB miss. 3. Check for unaligned, IO accesses 4. Victim TLB lookup. 5. If victim TLB misses, TLB refill 6. Do the memory access. 7. Return to code cache The advantage is that victim TLB can offer more associativity to a directly mapped TLB and thus potentially fewer page table walks while still keeping the time taken to flush within reasonable limits. However, placing a victim TLB before the refill path increase TLB refill path as the victim TLB is consulted before the TLB refill. The performance results demonstrate that the pros outweigh the cons. some performance results taken on SPECINT2006 train datasets and kernel boot and qemu configure script on an Intel(R) Xeon(R) CPU E5620 @ 2.40GHz Linux machine are shown in the Google Doc link below. https://docs.google.com/spreadsheets/d/1eiItzekZwNQOal_h-5iJmC4tMDi051m9qidi5_nwvH4/edit?usp=sharing In summary, victim TLB improves the performance of qemu-system-x86_64 by 11% on average on SPECINT2006, kernelboot and qemu configscript and with highest improvement of in 26% in 456.hmmer. And victim TLB does not result in any performance degradation in any of the measured benchmarks. Furthermore, the implemented victim TLB is architecture independent and is expected to benefit other architectures in QEMU as well. Although there are measurement fluctuations, the performance improvement is very significant and by no means in the range of noises. Signed-off-by: Xin Tong <trent.tong@gmail.com> Message-id: 1407202523-23553-1-git-send-email-trent.tong@gmail.com Reviewed-by: Peter Maydell <peter.maydell@linaro.org> Signed-off-by: Peter Maydell <peter.maydell@linaro.org>
2014-08-05 01:35:23 +00:00
memset(env->tlb_v_table, -1, sizeof(env->tlb_v_table));
tcg: consistently access cpu->tb_jmp_cache atomically Some code paths can lead to atomic accesses racing with memset() on cpu->tb_jmp_cache, which can result in torn reads/writes and is undefined behaviour in C11. These torn accesses are unlikely to show up as bugs, but from code inspection they seem possible. For example, tb_phys_invalidate does: /* remove the TB from the hash list */ h = tb_jmp_cache_hash_func(tb->pc); CPU_FOREACH(cpu) { if (atomic_read(&cpu->tb_jmp_cache[h]) == tb) { atomic_set(&cpu->tb_jmp_cache[h], NULL); } } Here atomic_set might race with a concurrent memset (such as the ones scheduled via "unsafe" async work, e.g. tlb_flush_page) and therefore we might end up with a torn pointer (or who knows what, because we are under undefined behaviour). This patch converts parallel accesses to cpu->tb_jmp_cache to use atomic primitives, thereby bringing these accesses back to defined behaviour. The price to pay is to potentially execute more instructions when clearing cpu->tb_jmp_cache, but given how infrequently they happen and the small size of the cache, the performance impact I have measured is within noise range when booting debian-arm. Note that under "safe async" work (e.g. do_tb_flush) we could use memset because no other vcpus are running. However I'm keeping these accesses atomic as well to keep things simple and to avoid confusing analysis tools such as ThreadSanitizer. Reviewed-by: Paolo Bonzini <pbonzini@redhat.com> Reviewed-by: Richard Henderson <rth@twiddle.net> Signed-off-by: Emilio G. Cota <cota@braap.org> Message-Id: <1497486973-25845-1-git-send-email-cota@braap.org> Signed-off-by: Richard Henderson <rth@twiddle.net>
2017-06-15 00:36:13 +00:00
cpu_tb_jmp_cache_clear(cpu);
implementing victim TLB for QEMU system emulated TLB QEMU system mode page table walks are expensive. Taken by running QEMU qemu-system-x86_64 system mode on Intel PIN , a TLB miss and walking a 4-level page tables in guest Linux OS takes ~450 X86 instructions on average. QEMU system mode TLB is implemented using a directly-mapped hashtable. This structure suffers from conflict misses. Increasing the associativity of the TLB may not be the solution to conflict misses as all the ways may have to be walked in serial. A victim TLB is a TLB used to hold translations evicted from the primary TLB upon replacement. The victim TLB lies between the main TLB and its refill path. Victim TLB is of greater associativity (fully associative in this patch). It takes longer to lookup the victim TLB, but its likely better than a full page table walk. The memory translation path is changed as follows : Before Victim TLB: 1. Inline TLB lookup 2. Exit code cache on TLB miss. 3. Check for unaligned, IO accesses 4. TLB refill. 5. Do the memory access. 6. Return to code cache. After Victim TLB: 1. Inline TLB lookup 2. Exit code cache on TLB miss. 3. Check for unaligned, IO accesses 4. Victim TLB lookup. 5. If victim TLB misses, TLB refill 6. Do the memory access. 7. Return to code cache The advantage is that victim TLB can offer more associativity to a directly mapped TLB and thus potentially fewer page table walks while still keeping the time taken to flush within reasonable limits. However, placing a victim TLB before the refill path increase TLB refill path as the victim TLB is consulted before the TLB refill. The performance results demonstrate that the pros outweigh the cons. some performance results taken on SPECINT2006 train datasets and kernel boot and qemu configure script on an Intel(R) Xeon(R) CPU E5620 @ 2.40GHz Linux machine are shown in the Google Doc link below. https://docs.google.com/spreadsheets/d/1eiItzekZwNQOal_h-5iJmC4tMDi051m9qidi5_nwvH4/edit?usp=sharing In summary, victim TLB improves the performance of qemu-system-x86_64 by 11% on average on SPECINT2006, kernelboot and qemu configscript and with highest improvement of in 26% in 456.hmmer. And victim TLB does not result in any performance degradation in any of the measured benchmarks. Furthermore, the implemented victim TLB is architecture independent and is expected to benefit other architectures in QEMU as well. Although there are measurement fluctuations, the performance improvement is very significant and by no means in the range of noises. Signed-off-by: Xin Tong <trent.tong@gmail.com> Message-id: 1407202523-23553-1-git-send-email-trent.tong@gmail.com Reviewed-by: Peter Maydell <peter.maydell@linaro.org> Signed-off-by: Peter Maydell <peter.maydell@linaro.org>
2014-08-05 01:35:23 +00:00
env->vtlb_index = 0;
env->tlb_flush_addr = -1;
env->tlb_flush_mask = 0;
tb_unlock();
atomic_mb_set(&cpu->pending_tlb_flush, 0);
}
static void tlb_flush_global_async_work(CPUState *cpu, run_on_cpu_data data)
{
tlb_flush_nocheck(cpu);
}
void tlb_flush(CPUState *cpu)
{
if (cpu->created && !qemu_cpu_is_self(cpu)) {
if (atomic_mb_read(&cpu->pending_tlb_flush) != ALL_MMUIDX_BITS) {
atomic_mb_set(&cpu->pending_tlb_flush, ALL_MMUIDX_BITS);
async_run_on_cpu(cpu, tlb_flush_global_async_work,
RUN_ON_CPU_NULL);
}
} else {
tlb_flush_nocheck(cpu);
}
}
void tlb_flush_all_cpus(CPUState *src_cpu)
{
const run_on_cpu_func fn = tlb_flush_global_async_work;
flush_all_helper(src_cpu, fn, RUN_ON_CPU_NULL);
fn(src_cpu, RUN_ON_CPU_NULL);
}
void tlb_flush_all_cpus_synced(CPUState *src_cpu)
{
const run_on_cpu_func fn = tlb_flush_global_async_work;
flush_all_helper(src_cpu, fn, RUN_ON_CPU_NULL);
async_safe_run_on_cpu(src_cpu, fn, RUN_ON_CPU_NULL);
}
static void tlb_flush_by_mmuidx_async_work(CPUState *cpu, run_on_cpu_data data)
{
CPUArchState *env = cpu->env_ptr;
unsigned long mmu_idx_bitmask = data.host_int;
int mmu_idx;
assert_cpu_is_self(cpu);
tb_lock();
tlb_debug("start: mmu_idx:0x%04lx\n", mmu_idx_bitmask);
for (mmu_idx = 0; mmu_idx < NB_MMU_MODES; mmu_idx++) {
if (test_bit(mmu_idx, &mmu_idx_bitmask)) {
tlb_debug("%d\n", mmu_idx);
memset(env->tlb_table[mmu_idx], -1, sizeof(env->tlb_table[0]));
memset(env->tlb_v_table[mmu_idx], -1, sizeof(env->tlb_v_table[0]));
}
}
tcg: consistently access cpu->tb_jmp_cache atomically Some code paths can lead to atomic accesses racing with memset() on cpu->tb_jmp_cache, which can result in torn reads/writes and is undefined behaviour in C11. These torn accesses are unlikely to show up as bugs, but from code inspection they seem possible. For example, tb_phys_invalidate does: /* remove the TB from the hash list */ h = tb_jmp_cache_hash_func(tb->pc); CPU_FOREACH(cpu) { if (atomic_read(&cpu->tb_jmp_cache[h]) == tb) { atomic_set(&cpu->tb_jmp_cache[h], NULL); } } Here atomic_set might race with a concurrent memset (such as the ones scheduled via "unsafe" async work, e.g. tlb_flush_page) and therefore we might end up with a torn pointer (or who knows what, because we are under undefined behaviour). This patch converts parallel accesses to cpu->tb_jmp_cache to use atomic primitives, thereby bringing these accesses back to defined behaviour. The price to pay is to potentially execute more instructions when clearing cpu->tb_jmp_cache, but given how infrequently they happen and the small size of the cache, the performance impact I have measured is within noise range when booting debian-arm. Note that under "safe async" work (e.g. do_tb_flush) we could use memset because no other vcpus are running. However I'm keeping these accesses atomic as well to keep things simple and to avoid confusing analysis tools such as ThreadSanitizer. Reviewed-by: Paolo Bonzini <pbonzini@redhat.com> Reviewed-by: Richard Henderson <rth@twiddle.net> Signed-off-by: Emilio G. Cota <cota@braap.org> Message-Id: <1497486973-25845-1-git-send-email-cota@braap.org> Signed-off-by: Richard Henderson <rth@twiddle.net>
2017-06-15 00:36:13 +00:00
cpu_tb_jmp_cache_clear(cpu);
tlb_debug("done\n");
tb_unlock();
}
void tlb_flush_by_mmuidx(CPUState *cpu, uint16_t idxmap)
{
tlb_debug("mmu_idx: 0x%" PRIx16 "\n", idxmap);
if (!qemu_cpu_is_self(cpu)) {
uint16_t pending_flushes = idxmap;
pending_flushes &= ~atomic_mb_read(&cpu->pending_tlb_flush);
if (pending_flushes) {
tlb_debug("reduced mmu_idx: 0x%" PRIx16 "\n", pending_flushes);
atomic_or(&cpu->pending_tlb_flush, pending_flushes);
async_run_on_cpu(cpu, tlb_flush_by_mmuidx_async_work,
RUN_ON_CPU_HOST_INT(pending_flushes));
}
} else {
tlb_flush_by_mmuidx_async_work(cpu,
RUN_ON_CPU_HOST_INT(idxmap));
}
}
void tlb_flush_by_mmuidx_all_cpus(CPUState *src_cpu, uint16_t idxmap)
{
const run_on_cpu_func fn = tlb_flush_by_mmuidx_async_work;
tlb_debug("mmu_idx: 0x%"PRIx16"\n", idxmap);
flush_all_helper(src_cpu, fn, RUN_ON_CPU_HOST_INT(idxmap));
fn(src_cpu, RUN_ON_CPU_HOST_INT(idxmap));
}
void tlb_flush_by_mmuidx_all_cpus_synced(CPUState *src_cpu,
uint16_t idxmap)
{
const run_on_cpu_func fn = tlb_flush_by_mmuidx_async_work;
tlb_debug("mmu_idx: 0x%"PRIx16"\n", idxmap);
flush_all_helper(src_cpu, fn, RUN_ON_CPU_HOST_INT(idxmap));
async_safe_run_on_cpu(src_cpu, fn, RUN_ON_CPU_HOST_INT(idxmap));
}
static inline void tlb_flush_entry(CPUTLBEntry *tlb_entry, target_ulong addr)
{
if (addr == (tlb_entry->addr_read &
(TARGET_PAGE_MASK | TLB_INVALID_MASK)) ||
addr == (tlb_entry->addr_write &
(TARGET_PAGE_MASK | TLB_INVALID_MASK)) ||
addr == (tlb_entry->addr_code &
(TARGET_PAGE_MASK | TLB_INVALID_MASK))) {
memset(tlb_entry, -1, sizeof(*tlb_entry));
}
}
static void tlb_flush_page_async_work(CPUState *cpu, run_on_cpu_data data)
{
CPUArchState *env = cpu->env_ptr;
target_ulong addr = (target_ulong) data.target_ptr;
int i;
int mmu_idx;
assert_cpu_is_self(cpu);
tlb_debug("page :" TARGET_FMT_lx "\n", addr);
/* Check if we need to flush due to large pages. */
if ((addr & env->tlb_flush_mask) == env->tlb_flush_addr) {
tlb_debug("forcing full flush ("
TARGET_FMT_lx "/" TARGET_FMT_lx ")\n",
env->tlb_flush_addr, env->tlb_flush_mask);
tlb_flush(cpu);
return;
}
addr &= TARGET_PAGE_MASK;
i = (addr >> TARGET_PAGE_BITS) & (CPU_TLB_SIZE - 1);
for (mmu_idx = 0; mmu_idx < NB_MMU_MODES; mmu_idx++) {
tlb_flush_entry(&env->tlb_table[mmu_idx][i], addr);
}
implementing victim TLB for QEMU system emulated TLB QEMU system mode page table walks are expensive. Taken by running QEMU qemu-system-x86_64 system mode on Intel PIN , a TLB miss and walking a 4-level page tables in guest Linux OS takes ~450 X86 instructions on average. QEMU system mode TLB is implemented using a directly-mapped hashtable. This structure suffers from conflict misses. Increasing the associativity of the TLB may not be the solution to conflict misses as all the ways may have to be walked in serial. A victim TLB is a TLB used to hold translations evicted from the primary TLB upon replacement. The victim TLB lies between the main TLB and its refill path. Victim TLB is of greater associativity (fully associative in this patch). It takes longer to lookup the victim TLB, but its likely better than a full page table walk. The memory translation path is changed as follows : Before Victim TLB: 1. Inline TLB lookup 2. Exit code cache on TLB miss. 3. Check for unaligned, IO accesses 4. TLB refill. 5. Do the memory access. 6. Return to code cache. After Victim TLB: 1. Inline TLB lookup 2. Exit code cache on TLB miss. 3. Check for unaligned, IO accesses 4. Victim TLB lookup. 5. If victim TLB misses, TLB refill 6. Do the memory access. 7. Return to code cache The advantage is that victim TLB can offer more associativity to a directly mapped TLB and thus potentially fewer page table walks while still keeping the time taken to flush within reasonable limits. However, placing a victim TLB before the refill path increase TLB refill path as the victim TLB is consulted before the TLB refill. The performance results demonstrate that the pros outweigh the cons. some performance results taken on SPECINT2006 train datasets and kernel boot and qemu configure script on an Intel(R) Xeon(R) CPU E5620 @ 2.40GHz Linux machine are shown in the Google Doc link below. https://docs.google.com/spreadsheets/d/1eiItzekZwNQOal_h-5iJmC4tMDi051m9qidi5_nwvH4/edit?usp=sharing In summary, victim TLB improves the performance of qemu-system-x86_64 by 11% on average on SPECINT2006, kernelboot and qemu configscript and with highest improvement of in 26% in 456.hmmer. And victim TLB does not result in any performance degradation in any of the measured benchmarks. Furthermore, the implemented victim TLB is architecture independent and is expected to benefit other architectures in QEMU as well. Although there are measurement fluctuations, the performance improvement is very significant and by no means in the range of noises. Signed-off-by: Xin Tong <trent.tong@gmail.com> Message-id: 1407202523-23553-1-git-send-email-trent.tong@gmail.com Reviewed-by: Peter Maydell <peter.maydell@linaro.org> Signed-off-by: Peter Maydell <peter.maydell@linaro.org>
2014-08-05 01:35:23 +00:00
/* check whether there are entries that need to be flushed in the vtlb */
for (mmu_idx = 0; mmu_idx < NB_MMU_MODES; mmu_idx++) {
int k;
for (k = 0; k < CPU_VTLB_SIZE; k++) {
tlb_flush_entry(&env->tlb_v_table[mmu_idx][k], addr);
}
}
tb_flush_jmp_cache(cpu, addr);
}
void tlb_flush_page(CPUState *cpu, target_ulong addr)
{
tlb_debug("page :" TARGET_FMT_lx "\n", addr);
if (!qemu_cpu_is_self(cpu)) {
async_run_on_cpu(cpu, tlb_flush_page_async_work,
RUN_ON_CPU_TARGET_PTR(addr));
} else {
tlb_flush_page_async_work(cpu, RUN_ON_CPU_TARGET_PTR(addr));
}
}
/* As we are going to hijack the bottom bits of the page address for a
* mmuidx bit mask we need to fail to build if we can't do that
*/
QEMU_BUILD_BUG_ON(NB_MMU_MODES > TARGET_PAGE_BITS_MIN);
static void tlb_flush_page_by_mmuidx_async_work(CPUState *cpu,
run_on_cpu_data data)
{
CPUArchState *env = cpu->env_ptr;
target_ulong addr_and_mmuidx = (target_ulong) data.target_ptr;
target_ulong addr = addr_and_mmuidx & TARGET_PAGE_MASK;
unsigned long mmu_idx_bitmap = addr_and_mmuidx & ALL_MMUIDX_BITS;
int page = (addr >> TARGET_PAGE_BITS) & (CPU_TLB_SIZE - 1);
int mmu_idx;
int i;
assert_cpu_is_self(cpu);
tlb_debug("page:%d addr:"TARGET_FMT_lx" mmu_idx:0x%lx\n",
page, addr, mmu_idx_bitmap);
for (mmu_idx = 0; mmu_idx < NB_MMU_MODES; mmu_idx++) {
if (test_bit(mmu_idx, &mmu_idx_bitmap)) {
tlb_flush_entry(&env->tlb_table[mmu_idx][page], addr);
/* check whether there are vltb entries that need to be flushed */
for (i = 0; i < CPU_VTLB_SIZE; i++) {
tlb_flush_entry(&env->tlb_v_table[mmu_idx][i], addr);
}
}
}
tb_flush_jmp_cache(cpu, addr);
}
static void tlb_check_page_and_flush_by_mmuidx_async_work(CPUState *cpu,
run_on_cpu_data data)
{
CPUArchState *env = cpu->env_ptr;
target_ulong addr_and_mmuidx = (target_ulong) data.target_ptr;
target_ulong addr = addr_and_mmuidx & TARGET_PAGE_MASK;
unsigned long mmu_idx_bitmap = addr_and_mmuidx & ALL_MMUIDX_BITS;
tlb_debug("addr:"TARGET_FMT_lx" mmu_idx: %04lx\n", addr, mmu_idx_bitmap);
/* Check if we need to flush due to large pages. */
if ((addr & env->tlb_flush_mask) == env->tlb_flush_addr) {
tlb_debug("forced full flush ("
TARGET_FMT_lx "/" TARGET_FMT_lx ")\n",
env->tlb_flush_addr, env->tlb_flush_mask);
tlb_flush_by_mmuidx_async_work(cpu,
RUN_ON_CPU_HOST_INT(mmu_idx_bitmap));
} else {
tlb_flush_page_by_mmuidx_async_work(cpu, data);
}
}
void tlb_flush_page_by_mmuidx(CPUState *cpu, target_ulong addr, uint16_t idxmap)
{
target_ulong addr_and_mmu_idx;
tlb_debug("addr: "TARGET_FMT_lx" mmu_idx:%" PRIx16 "\n", addr, idxmap);
/* This should already be page aligned */
addr_and_mmu_idx = addr & TARGET_PAGE_MASK;
addr_and_mmu_idx |= idxmap;
if (!qemu_cpu_is_self(cpu)) {
async_run_on_cpu(cpu, tlb_check_page_and_flush_by_mmuidx_async_work,
RUN_ON_CPU_TARGET_PTR(addr_and_mmu_idx));
} else {
tlb_check_page_and_flush_by_mmuidx_async_work(
cpu, RUN_ON_CPU_TARGET_PTR(addr_and_mmu_idx));
}
}
void tlb_flush_page_by_mmuidx_all_cpus(CPUState *src_cpu, target_ulong addr,
uint16_t idxmap)
{
const run_on_cpu_func fn = tlb_check_page_and_flush_by_mmuidx_async_work;
target_ulong addr_and_mmu_idx;
tlb_debug("addr: "TARGET_FMT_lx" mmu_idx:%"PRIx16"\n", addr, idxmap);
/* This should already be page aligned */
addr_and_mmu_idx = addr & TARGET_PAGE_MASK;
addr_and_mmu_idx |= idxmap;
flush_all_helper(src_cpu, fn, RUN_ON_CPU_TARGET_PTR(addr_and_mmu_idx));
fn(src_cpu, RUN_ON_CPU_TARGET_PTR(addr_and_mmu_idx));
}
void tlb_flush_page_by_mmuidx_all_cpus_synced(CPUState *src_cpu,
target_ulong addr,
uint16_t idxmap)
{
const run_on_cpu_func fn = tlb_check_page_and_flush_by_mmuidx_async_work;
target_ulong addr_and_mmu_idx;
tlb_debug("addr: "TARGET_FMT_lx" mmu_idx:%"PRIx16"\n", addr, idxmap);
/* This should already be page aligned */
addr_and_mmu_idx = addr & TARGET_PAGE_MASK;
addr_and_mmu_idx |= idxmap;
flush_all_helper(src_cpu, fn, RUN_ON_CPU_TARGET_PTR(addr_and_mmu_idx));
async_safe_run_on_cpu(src_cpu, fn, RUN_ON_CPU_TARGET_PTR(addr_and_mmu_idx));
}
void tlb_flush_page_all_cpus(CPUState *src, target_ulong addr)
{
const run_on_cpu_func fn = tlb_flush_page_async_work;
flush_all_helper(src, fn, RUN_ON_CPU_TARGET_PTR(addr));
fn(src, RUN_ON_CPU_TARGET_PTR(addr));
}
void tlb_flush_page_all_cpus_synced(CPUState *src,
target_ulong addr)
{
const run_on_cpu_func fn = tlb_flush_page_async_work;
flush_all_helper(src, fn, RUN_ON_CPU_TARGET_PTR(addr));
async_safe_run_on_cpu(src, fn, RUN_ON_CPU_TARGET_PTR(addr));
}
/* update the TLBs so that writes to code in the virtual page 'addr'
can be detected */
void tlb_protect_code(ram_addr_t ram_addr)
{
cpu_physical_memory_test_and_clear_dirty(ram_addr, TARGET_PAGE_SIZE,
DIRTY_MEMORY_CODE);
}
/* update the TLB so that writes in physical page 'phys_addr' are no longer
tested for self modifying code */
void tlb_unprotect_code(ram_addr_t ram_addr)
{
cpu_physical_memory_set_dirty_flag(ram_addr, DIRTY_MEMORY_CODE);
}
/*
* Dirty write flag handling
*
* When the TCG code writes to a location it looks up the address in
* the TLB and uses that data to compute the final address. If any of
* the lower bits of the address are set then the slow path is forced.
* There are a number of reasons to do this but for normal RAM the
* most usual is detecting writes to code regions which may invalidate
* generated code.
*
* Because we want other vCPUs to respond to changes straight away we
* update the te->addr_write field atomically. If the TLB entry has
* been changed by the vCPU in the mean time we skip the update.
*
* As this function uses atomic accesses we also need to ensure
* updates to tlb_entries follow the same access rules. We don't need
* to worry about this for oversized guests as MTTCG is disabled for
* them.
*/
static void tlb_reset_dirty_range(CPUTLBEntry *tlb_entry, uintptr_t start,
uintptr_t length)
{
#if TCG_OVERSIZED_GUEST
uintptr_t addr = tlb_entry->addr_write;
if ((addr & (TLB_INVALID_MASK | TLB_MMIO | TLB_NOTDIRTY)) == 0) {
addr &= TARGET_PAGE_MASK;
addr += tlb_entry->addend;
if ((addr - start) < length) {
tlb_entry->addr_write |= TLB_NOTDIRTY;
}
}
#else
/* paired with atomic_mb_set in tlb_set_page_with_attrs */
uintptr_t orig_addr = atomic_mb_read(&tlb_entry->addr_write);
uintptr_t addr = orig_addr;
if ((addr & (TLB_INVALID_MASK | TLB_MMIO | TLB_NOTDIRTY)) == 0) {
addr &= TARGET_PAGE_MASK;
addr += atomic_read(&tlb_entry->addend);
if ((addr - start) < length) {
uintptr_t notdirty_addr = orig_addr | TLB_NOTDIRTY;
atomic_cmpxchg(&tlb_entry->addr_write, orig_addr, notdirty_addr);
}
}
#endif
}
/* For atomic correctness when running MTTCG we need to use the right
* primitives when copying entries */
static inline void copy_tlb_helper(CPUTLBEntry *d, CPUTLBEntry *s,
bool atomic_set)
{
#if TCG_OVERSIZED_GUEST
*d = *s;
#else
if (atomic_set) {
d->addr_read = s->addr_read;
d->addr_code = s->addr_code;
atomic_set(&d->addend, atomic_read(&s->addend));
/* Pairs with flag setting in tlb_reset_dirty_range */
atomic_mb_set(&d->addr_write, atomic_read(&s->addr_write));
} else {
d->addr_read = s->addr_read;
d->addr_write = atomic_read(&s->addr_write);
d->addr_code = s->addr_code;
d->addend = atomic_read(&s->addend);
}
#endif
}
/* This is a cross vCPU call (i.e. another vCPU resetting the flags of
* the target vCPU). As such care needs to be taken that we don't
* dangerously race with another vCPU update. The only thing actually
* updated is the target TLB entry ->addr_write flags.
*/
void tlb_reset_dirty(CPUState *cpu, ram_addr_t start1, ram_addr_t length)
{
CPUArchState *env;
int mmu_idx;
env = cpu->env_ptr;
for (mmu_idx = 0; mmu_idx < NB_MMU_MODES; mmu_idx++) {
unsigned int i;
for (i = 0; i < CPU_TLB_SIZE; i++) {
tlb_reset_dirty_range(&env->tlb_table[mmu_idx][i],
start1, length);
}
implementing victim TLB for QEMU system emulated TLB QEMU system mode page table walks are expensive. Taken by running QEMU qemu-system-x86_64 system mode on Intel PIN , a TLB miss and walking a 4-level page tables in guest Linux OS takes ~450 X86 instructions on average. QEMU system mode TLB is implemented using a directly-mapped hashtable. This structure suffers from conflict misses. Increasing the associativity of the TLB may not be the solution to conflict misses as all the ways may have to be walked in serial. A victim TLB is a TLB used to hold translations evicted from the primary TLB upon replacement. The victim TLB lies between the main TLB and its refill path. Victim TLB is of greater associativity (fully associative in this patch). It takes longer to lookup the victim TLB, but its likely better than a full page table walk. The memory translation path is changed as follows : Before Victim TLB: 1. Inline TLB lookup 2. Exit code cache on TLB miss. 3. Check for unaligned, IO accesses 4. TLB refill. 5. Do the memory access. 6. Return to code cache. After Victim TLB: 1. Inline TLB lookup 2. Exit code cache on TLB miss. 3. Check for unaligned, IO accesses 4. Victim TLB lookup. 5. If victim TLB misses, TLB refill 6. Do the memory access. 7. Return to code cache The advantage is that victim TLB can offer more associativity to a directly mapped TLB and thus potentially fewer page table walks while still keeping the time taken to flush within reasonable limits. However, placing a victim TLB before the refill path increase TLB refill path as the victim TLB is consulted before the TLB refill. The performance results demonstrate that the pros outweigh the cons. some performance results taken on SPECINT2006 train datasets and kernel boot and qemu configure script on an Intel(R) Xeon(R) CPU E5620 @ 2.40GHz Linux machine are shown in the Google Doc link below. https://docs.google.com/spreadsheets/d/1eiItzekZwNQOal_h-5iJmC4tMDi051m9qidi5_nwvH4/edit?usp=sharing In summary, victim TLB improves the performance of qemu-system-x86_64 by 11% on average on SPECINT2006, kernelboot and qemu configscript and with highest improvement of in 26% in 456.hmmer. And victim TLB does not result in any performance degradation in any of the measured benchmarks. Furthermore, the implemented victim TLB is architecture independent and is expected to benefit other architectures in QEMU as well. Although there are measurement fluctuations, the performance improvement is very significant and by no means in the range of noises. Signed-off-by: Xin Tong <trent.tong@gmail.com> Message-id: 1407202523-23553-1-git-send-email-trent.tong@gmail.com Reviewed-by: Peter Maydell <peter.maydell@linaro.org> Signed-off-by: Peter Maydell <peter.maydell@linaro.org>
2014-08-05 01:35:23 +00:00
for (i = 0; i < CPU_VTLB_SIZE; i++) {
tlb_reset_dirty_range(&env->tlb_v_table[mmu_idx][i],
start1, length);
}
}
}
static inline void tlb_set_dirty1(CPUTLBEntry *tlb_entry, target_ulong vaddr)
{
if (tlb_entry->addr_write == (vaddr | TLB_NOTDIRTY)) {
tlb_entry->addr_write = vaddr;
}
}
/* update the TLB corresponding to virtual page vaddr
so that it is no longer dirty */
void tlb_set_dirty(CPUState *cpu, target_ulong vaddr)
{
CPUArchState *env = cpu->env_ptr;
int i;
int mmu_idx;
assert_cpu_is_self(cpu);
vaddr &= TARGET_PAGE_MASK;
i = (vaddr >> TARGET_PAGE_BITS) & (CPU_TLB_SIZE - 1);
for (mmu_idx = 0; mmu_idx < NB_MMU_MODES; mmu_idx++) {
tlb_set_dirty1(&env->tlb_table[mmu_idx][i], vaddr);
}
implementing victim TLB for QEMU system emulated TLB QEMU system mode page table walks are expensive. Taken by running QEMU qemu-system-x86_64 system mode on Intel PIN , a TLB miss and walking a 4-level page tables in guest Linux OS takes ~450 X86 instructions on average. QEMU system mode TLB is implemented using a directly-mapped hashtable. This structure suffers from conflict misses. Increasing the associativity of the TLB may not be the solution to conflict misses as all the ways may have to be walked in serial. A victim TLB is a TLB used to hold translations evicted from the primary TLB upon replacement. The victim TLB lies between the main TLB and its refill path. Victim TLB is of greater associativity (fully associative in this patch). It takes longer to lookup the victim TLB, but its likely better than a full page table walk. The memory translation path is changed as follows : Before Victim TLB: 1. Inline TLB lookup 2. Exit code cache on TLB miss. 3. Check for unaligned, IO accesses 4. TLB refill. 5. Do the memory access. 6. Return to code cache. After Victim TLB: 1. Inline TLB lookup 2. Exit code cache on TLB miss. 3. Check for unaligned, IO accesses 4. Victim TLB lookup. 5. If victim TLB misses, TLB refill 6. Do the memory access. 7. Return to code cache The advantage is that victim TLB can offer more associativity to a directly mapped TLB and thus potentially fewer page table walks while still keeping the time taken to flush within reasonable limits. However, placing a victim TLB before the refill path increase TLB refill path as the victim TLB is consulted before the TLB refill. The performance results demonstrate that the pros outweigh the cons. some performance results taken on SPECINT2006 train datasets and kernel boot and qemu configure script on an Intel(R) Xeon(R) CPU E5620 @ 2.40GHz Linux machine are shown in the Google Doc link below. https://docs.google.com/spreadsheets/d/1eiItzekZwNQOal_h-5iJmC4tMDi051m9qidi5_nwvH4/edit?usp=sharing In summary, victim TLB improves the performance of qemu-system-x86_64 by 11% on average on SPECINT2006, kernelboot and qemu configscript and with highest improvement of in 26% in 456.hmmer. And victim TLB does not result in any performance degradation in any of the measured benchmarks. Furthermore, the implemented victim TLB is architecture independent and is expected to benefit other architectures in QEMU as well. Although there are measurement fluctuations, the performance improvement is very significant and by no means in the range of noises. Signed-off-by: Xin Tong <trent.tong@gmail.com> Message-id: 1407202523-23553-1-git-send-email-trent.tong@gmail.com Reviewed-by: Peter Maydell <peter.maydell@linaro.org> Signed-off-by: Peter Maydell <peter.maydell@linaro.org>
2014-08-05 01:35:23 +00:00
for (mmu_idx = 0; mmu_idx < NB_MMU_MODES; mmu_idx++) {
int k;
for (k = 0; k < CPU_VTLB_SIZE; k++) {
tlb_set_dirty1(&env->tlb_v_table[mmu_idx][k], vaddr);
}
}
}
/* Our TLB does not support large pages, so remember the area covered by
large pages and trigger a full TLB flush if these are invalidated. */
static void tlb_add_large_page(CPUArchState *env, target_ulong vaddr,
target_ulong size)
{
target_ulong mask = ~(size - 1);
if (env->tlb_flush_addr == (target_ulong)-1) {
env->tlb_flush_addr = vaddr & mask;
env->tlb_flush_mask = mask;
return;
}
/* Extend the existing region to include the new page.
This is a compromise between unnecessary flushes and the cost
of maintaining a full variable size TLB. */
mask &= env->tlb_flush_mask;
while (((env->tlb_flush_addr ^ vaddr) & mask) != 0) {
mask <<= 1;
}
env->tlb_flush_addr &= mask;
env->tlb_flush_mask = mask;
}
/* Add a new TLB entry. At most one entry for a given virtual address
* is permitted. Only a single TARGET_PAGE_SIZE region is mapped, the
* supplied size is only used by tlb_flush_page.
*
* Called from TCG-generated code, which is under an RCU read-side
* critical section.
*/
void tlb_set_page_with_attrs(CPUState *cpu, target_ulong vaddr,
hwaddr paddr, MemTxAttrs attrs, int prot,
int mmu_idx, target_ulong size)
{
CPUArchState *env = cpu->env_ptr;
MemoryRegionSection *section;
unsigned int index;
target_ulong address;
target_ulong code_address;
uintptr_t addend;
CPUTLBEntry *te, *tv, tn;
hwaddr iotlb, xlat, sz;
implementing victim TLB for QEMU system emulated TLB QEMU system mode page table walks are expensive. Taken by running QEMU qemu-system-x86_64 system mode on Intel PIN , a TLB miss and walking a 4-level page tables in guest Linux OS takes ~450 X86 instructions on average. QEMU system mode TLB is implemented using a directly-mapped hashtable. This structure suffers from conflict misses. Increasing the associativity of the TLB may not be the solution to conflict misses as all the ways may have to be walked in serial. A victim TLB is a TLB used to hold translations evicted from the primary TLB upon replacement. The victim TLB lies between the main TLB and its refill path. Victim TLB is of greater associativity (fully associative in this patch). It takes longer to lookup the victim TLB, but its likely better than a full page table walk. The memory translation path is changed as follows : Before Victim TLB: 1. Inline TLB lookup 2. Exit code cache on TLB miss. 3. Check for unaligned, IO accesses 4. TLB refill. 5. Do the memory access. 6. Return to code cache. After Victim TLB: 1. Inline TLB lookup 2. Exit code cache on TLB miss. 3. Check for unaligned, IO accesses 4. Victim TLB lookup. 5. If victim TLB misses, TLB refill 6. Do the memory access. 7. Return to code cache The advantage is that victim TLB can offer more associativity to a directly mapped TLB and thus potentially fewer page table walks while still keeping the time taken to flush within reasonable limits. However, placing a victim TLB before the refill path increase TLB refill path as the victim TLB is consulted before the TLB refill. The performance results demonstrate that the pros outweigh the cons. some performance results taken on SPECINT2006 train datasets and kernel boot and qemu configure script on an Intel(R) Xeon(R) CPU E5620 @ 2.40GHz Linux machine are shown in the Google Doc link below. https://docs.google.com/spreadsheets/d/1eiItzekZwNQOal_h-5iJmC4tMDi051m9qidi5_nwvH4/edit?usp=sharing In summary, victim TLB improves the performance of qemu-system-x86_64 by 11% on average on SPECINT2006, kernelboot and qemu configscript and with highest improvement of in 26% in 456.hmmer. And victim TLB does not result in any performance degradation in any of the measured benchmarks. Furthermore, the implemented victim TLB is architecture independent and is expected to benefit other architectures in QEMU as well. Although there are measurement fluctuations, the performance improvement is very significant and by no means in the range of noises. Signed-off-by: Xin Tong <trent.tong@gmail.com> Message-id: 1407202523-23553-1-git-send-email-trent.tong@gmail.com Reviewed-by: Peter Maydell <peter.maydell@linaro.org> Signed-off-by: Peter Maydell <peter.maydell@linaro.org>
2014-08-05 01:35:23 +00:00
unsigned vidx = env->vtlb_index++ % CPU_VTLB_SIZE;
int asidx = cpu_asidx_from_attrs(cpu, attrs);
assert_cpu_is_self(cpu);
assert(size >= TARGET_PAGE_SIZE);
if (size != TARGET_PAGE_SIZE) {
tlb_add_large_page(env, vaddr, size);
}
sz = size;
section = address_space_translate_for_iotlb(cpu, asidx, paddr, &xlat, &sz);
assert(sz >= TARGET_PAGE_SIZE);
tlb_debug("vaddr=" TARGET_FMT_lx " paddr=0x" TARGET_FMT_plx
" prot=%x idx=%d\n",
vaddr, paddr, prot, mmu_idx);
address = vaddr;
if (!memory_region_is_ram(section->mr) && !memory_region_is_romd(section->mr)) {
/* IO memory case */
address |= TLB_MMIO;
addend = 0;
} else {
/* TLB_MMIO for rom/romd handled below */
addend = (uintptr_t)memory_region_get_ram_ptr(section->mr) + xlat;
}
code_address = address;
iotlb = memory_region_section_get_iotlb(cpu, section, vaddr, paddr, xlat,
prot, &address);
index = (vaddr >> TARGET_PAGE_BITS) & (CPU_TLB_SIZE - 1);
te = &env->tlb_table[mmu_idx][index];
implementing victim TLB for QEMU system emulated TLB QEMU system mode page table walks are expensive. Taken by running QEMU qemu-system-x86_64 system mode on Intel PIN , a TLB miss and walking a 4-level page tables in guest Linux OS takes ~450 X86 instructions on average. QEMU system mode TLB is implemented using a directly-mapped hashtable. This structure suffers from conflict misses. Increasing the associativity of the TLB may not be the solution to conflict misses as all the ways may have to be walked in serial. A victim TLB is a TLB used to hold translations evicted from the primary TLB upon replacement. The victim TLB lies between the main TLB and its refill path. Victim TLB is of greater associativity (fully associative in this patch). It takes longer to lookup the victim TLB, but its likely better than a full page table walk. The memory translation path is changed as follows : Before Victim TLB: 1. Inline TLB lookup 2. Exit code cache on TLB miss. 3. Check for unaligned, IO accesses 4. TLB refill. 5. Do the memory access. 6. Return to code cache. After Victim TLB: 1. Inline TLB lookup 2. Exit code cache on TLB miss. 3. Check for unaligned, IO accesses 4. Victim TLB lookup. 5. If victim TLB misses, TLB refill 6. Do the memory access. 7. Return to code cache The advantage is that victim TLB can offer more associativity to a directly mapped TLB and thus potentially fewer page table walks while still keeping the time taken to flush within reasonable limits. However, placing a victim TLB before the refill path increase TLB refill path as the victim TLB is consulted before the TLB refill. The performance results demonstrate that the pros outweigh the cons. some performance results taken on SPECINT2006 train datasets and kernel boot and qemu configure script on an Intel(R) Xeon(R) CPU E5620 @ 2.40GHz Linux machine are shown in the Google Doc link below. https://docs.google.com/spreadsheets/d/1eiItzekZwNQOal_h-5iJmC4tMDi051m9qidi5_nwvH4/edit?usp=sharing In summary, victim TLB improves the performance of qemu-system-x86_64 by 11% on average on SPECINT2006, kernelboot and qemu configscript and with highest improvement of in 26% in 456.hmmer. And victim TLB does not result in any performance degradation in any of the measured benchmarks. Furthermore, the implemented victim TLB is architecture independent and is expected to benefit other architectures in QEMU as well. Although there are measurement fluctuations, the performance improvement is very significant and by no means in the range of noises. Signed-off-by: Xin Tong <trent.tong@gmail.com> Message-id: 1407202523-23553-1-git-send-email-trent.tong@gmail.com Reviewed-by: Peter Maydell <peter.maydell@linaro.org> Signed-off-by: Peter Maydell <peter.maydell@linaro.org>
2014-08-05 01:35:23 +00:00
/* do not discard the translation in te, evict it into a victim tlb */
tv = &env->tlb_v_table[mmu_idx][vidx];
/* addr_write can race with tlb_reset_dirty_range */
copy_tlb_helper(tv, te, true);
implementing victim TLB for QEMU system emulated TLB QEMU system mode page table walks are expensive. Taken by running QEMU qemu-system-x86_64 system mode on Intel PIN , a TLB miss and walking a 4-level page tables in guest Linux OS takes ~450 X86 instructions on average. QEMU system mode TLB is implemented using a directly-mapped hashtable. This structure suffers from conflict misses. Increasing the associativity of the TLB may not be the solution to conflict misses as all the ways may have to be walked in serial. A victim TLB is a TLB used to hold translations evicted from the primary TLB upon replacement. The victim TLB lies between the main TLB and its refill path. Victim TLB is of greater associativity (fully associative in this patch). It takes longer to lookup the victim TLB, but its likely better than a full page table walk. The memory translation path is changed as follows : Before Victim TLB: 1. Inline TLB lookup 2. Exit code cache on TLB miss. 3. Check for unaligned, IO accesses 4. TLB refill. 5. Do the memory access. 6. Return to code cache. After Victim TLB: 1. Inline TLB lookup 2. Exit code cache on TLB miss. 3. Check for unaligned, IO accesses 4. Victim TLB lookup. 5. If victim TLB misses, TLB refill 6. Do the memory access. 7. Return to code cache The advantage is that victim TLB can offer more associativity to a directly mapped TLB and thus potentially fewer page table walks while still keeping the time taken to flush within reasonable limits. However, placing a victim TLB before the refill path increase TLB refill path as the victim TLB is consulted before the TLB refill. The performance results demonstrate that the pros outweigh the cons. some performance results taken on SPECINT2006 train datasets and kernel boot and qemu configure script on an Intel(R) Xeon(R) CPU E5620 @ 2.40GHz Linux machine are shown in the Google Doc link below. https://docs.google.com/spreadsheets/d/1eiItzekZwNQOal_h-5iJmC4tMDi051m9qidi5_nwvH4/edit?usp=sharing In summary, victim TLB improves the performance of qemu-system-x86_64 by 11% on average on SPECINT2006, kernelboot and qemu configscript and with highest improvement of in 26% in 456.hmmer. And victim TLB does not result in any performance degradation in any of the measured benchmarks. Furthermore, the implemented victim TLB is architecture independent and is expected to benefit other architectures in QEMU as well. Although there are measurement fluctuations, the performance improvement is very significant and by no means in the range of noises. Signed-off-by: Xin Tong <trent.tong@gmail.com> Message-id: 1407202523-23553-1-git-send-email-trent.tong@gmail.com Reviewed-by: Peter Maydell <peter.maydell@linaro.org> Signed-off-by: Peter Maydell <peter.maydell@linaro.org>
2014-08-05 01:35:23 +00:00
env->iotlb_v[mmu_idx][vidx] = env->iotlb[mmu_idx][index];
/* refill the tlb */
env->iotlb[mmu_idx][index].addr = iotlb - vaddr;
env->iotlb[mmu_idx][index].attrs = attrs;
/* Now calculate the new entry */
tn.addend = addend - vaddr;
if (prot & PAGE_READ) {
tn.addr_read = address;
} else {
tn.addr_read = -1;
}
if (prot & PAGE_EXEC) {
tn.addr_code = code_address;
} else {
tn.addr_code = -1;
}
tn.addr_write = -1;
if (prot & PAGE_WRITE) {
if ((memory_region_is_ram(section->mr) && section->readonly)
|| memory_region_is_romd(section->mr)) {
/* Write access calls the I/O callback. */
tn.addr_write = address | TLB_MMIO;
} else if (memory_region_is_ram(section->mr)
&& cpu_physical_memory_is_clean(
memory_region_get_ram_addr(section->mr) + xlat)) {
tn.addr_write = address | TLB_NOTDIRTY;
} else {
tn.addr_write = address;
}
}
/* Pairs with flag setting in tlb_reset_dirty_range */
copy_tlb_helper(te, &tn, true);
/* atomic_mb_set(&te->addr_write, write_address); */
}
/* Add a new TLB entry, but without specifying the memory
* transaction attributes to be used.
*/
void tlb_set_page(CPUState *cpu, target_ulong vaddr,
hwaddr paddr, int prot,
int mmu_idx, target_ulong size)
{
tlb_set_page_with_attrs(cpu, vaddr, paddr, MEMTXATTRS_UNSPECIFIED,
prot, mmu_idx, size);
}
static void report_bad_exec(CPUState *cpu, target_ulong addr)
{
/* Accidentally executing outside RAM or ROM is quite common for
* several user-error situations, so report it in a way that
* makes it clear that this isn't a QEMU bug and provide suggestions
* about what a user could do to fix things.
*/
error_report("Trying to execute code outside RAM or ROM at 0x"
TARGET_FMT_lx, addr);
error_printf("This usually means one of the following happened:\n\n"
"(1) You told QEMU to execute a kernel for the wrong machine "
"type, and it crashed on startup (eg trying to run a "
"raspberry pi kernel on a versatilepb QEMU machine)\n"
"(2) You didn't give QEMU a kernel or BIOS filename at all, "
"and QEMU executed a ROM full of no-op instructions until "
"it fell off the end\n"
"(3) Your guest kernel has a bug and crashed by jumping "
"off into nowhere\n\n"
"This is almost always one of the first two, so check your "
"command line and that you are using the right type of kernel "
"for this machine.\n"
"If you think option (3) is likely then you can try debugging "
"your guest with the -d debug options; in particular "
"-d guest_errors will cause the log to include a dump of the "
"guest register state at this point.\n\n"
"Execution cannot continue; stopping here.\n\n");
/* Report also to the logs, with more detail including register dump */
qemu_log_mask(LOG_GUEST_ERROR, "qemu: fatal: Trying to execute code "
"outside RAM or ROM at 0x" TARGET_FMT_lx "\n", addr);
log_cpu_state_mask(LOG_GUEST_ERROR, cpu, CPU_DUMP_FPU | CPU_DUMP_CCOP);
}
static inline ram_addr_t qemu_ram_addr_from_host_nofail(void *ptr)
{
ram_addr_t ram_addr;
ram_addr = qemu_ram_addr_from_host(ptr);
if (ram_addr == RAM_ADDR_INVALID) {
error_report("Bad ram pointer %p", ptr);
abort();
}
return ram_addr;
}
static uint64_t io_readx(CPUArchState *env, CPUIOTLBEntry *iotlbentry,
int mmu_idx,
target_ulong addr, uintptr_t retaddr, int size)
{
CPUState *cpu = ENV_GET_CPU(env);
hwaddr physaddr = iotlbentry->addr;
MemoryRegion *mr = iotlb_to_region(cpu, physaddr, iotlbentry->attrs);
uint64_t val;
tcg: drop global lock during TCG code execution This finally allows TCG to benefit from the iothread introduction: Drop the global mutex while running pure TCG CPU code. Reacquire the lock when entering MMIO or PIO emulation, or when leaving the TCG loop. We have to revert a few optimization for the current TCG threading model, namely kicking the TCG thread in qemu_mutex_lock_iothread and not kicking it in qemu_cpu_kick. We also need to disable RAM block reordering until we have a more efficient locking mechanism at hand. Still, a Linux x86 UP guest and my Musicpal ARM model boot fine here. These numbers demonstrate where we gain something: 20338 jan 20 0 331m 75m 6904 R 99 0.9 0:50.95 qemu-system-arm 20337 jan 20 0 331m 75m 6904 S 20 0.9 0:26.50 qemu-system-arm The guest CPU was fully loaded, but the iothread could still run mostly independent on a second core. Without the patch we don't get beyond 32206 jan 20 0 330m 73m 7036 R 82 0.9 1:06.00 qemu-system-arm 32204 jan 20 0 330m 73m 7036 S 21 0.9 0:17.03 qemu-system-arm We don't benefit significantly, though, when the guest is not fully loading a host CPU. Signed-off-by: Jan Kiszka <jan.kiszka@siemens.com> Message-Id: <1439220437-23957-10-git-send-email-fred.konrad@greensocs.com> [FK: Rebase, fix qemu_devices_reset deadlock, rm address_space_* mutex] Signed-off-by: KONRAD Frederic <fred.konrad@greensocs.com> [EGC: fixed iothread lock for cpu-exec IRQ handling] Signed-off-by: Emilio G. Cota <cota@braap.org> [AJB: -smp single-threaded fix, clean commit msg, BQL fixes] Signed-off-by: Alex Bennée <alex.bennee@linaro.org> Reviewed-by: Richard Henderson <rth@twiddle.net> Reviewed-by: Pranith Kumar <bobby.prani@gmail.com> [PM: target-arm changes] Acked-by: Peter Maydell <peter.maydell@linaro.org>
2017-02-23 18:29:11 +00:00
bool locked = false;
MemTxResult r;
physaddr = (physaddr & TARGET_PAGE_MASK) + addr;
cpu->mem_io_pc = retaddr;
if (mr != &io_mem_rom && mr != &io_mem_notdirty && !cpu->can_do_io) {
cpu_io_recompile(cpu, retaddr);
}
cpu->mem_io_vaddr = addr;
tcg: drop global lock during TCG code execution This finally allows TCG to benefit from the iothread introduction: Drop the global mutex while running pure TCG CPU code. Reacquire the lock when entering MMIO or PIO emulation, or when leaving the TCG loop. We have to revert a few optimization for the current TCG threading model, namely kicking the TCG thread in qemu_mutex_lock_iothread and not kicking it in qemu_cpu_kick. We also need to disable RAM block reordering until we have a more efficient locking mechanism at hand. Still, a Linux x86 UP guest and my Musicpal ARM model boot fine here. These numbers demonstrate where we gain something: 20338 jan 20 0 331m 75m 6904 R 99 0.9 0:50.95 qemu-system-arm 20337 jan 20 0 331m 75m 6904 S 20 0.9 0:26.50 qemu-system-arm The guest CPU was fully loaded, but the iothread could still run mostly independent on a second core. Without the patch we don't get beyond 32206 jan 20 0 330m 73m 7036 R 82 0.9 1:06.00 qemu-system-arm 32204 jan 20 0 330m 73m 7036 S 21 0.9 0:17.03 qemu-system-arm We don't benefit significantly, though, when the guest is not fully loading a host CPU. Signed-off-by: Jan Kiszka <jan.kiszka@siemens.com> Message-Id: <1439220437-23957-10-git-send-email-fred.konrad@greensocs.com> [FK: Rebase, fix qemu_devices_reset deadlock, rm address_space_* mutex] Signed-off-by: KONRAD Frederic <fred.konrad@greensocs.com> [EGC: fixed iothread lock for cpu-exec IRQ handling] Signed-off-by: Emilio G. Cota <cota@braap.org> [AJB: -smp single-threaded fix, clean commit msg, BQL fixes] Signed-off-by: Alex Bennée <alex.bennee@linaro.org> Reviewed-by: Richard Henderson <rth@twiddle.net> Reviewed-by: Pranith Kumar <bobby.prani@gmail.com> [PM: target-arm changes] Acked-by: Peter Maydell <peter.maydell@linaro.org>
2017-02-23 18:29:11 +00:00
if (mr->global_locking && !qemu_mutex_iothread_locked()) {
tcg: drop global lock during TCG code execution This finally allows TCG to benefit from the iothread introduction: Drop the global mutex while running pure TCG CPU code. Reacquire the lock when entering MMIO or PIO emulation, or when leaving the TCG loop. We have to revert a few optimization for the current TCG threading model, namely kicking the TCG thread in qemu_mutex_lock_iothread and not kicking it in qemu_cpu_kick. We also need to disable RAM block reordering until we have a more efficient locking mechanism at hand. Still, a Linux x86 UP guest and my Musicpal ARM model boot fine here. These numbers demonstrate where we gain something: 20338 jan 20 0 331m 75m 6904 R 99 0.9 0:50.95 qemu-system-arm 20337 jan 20 0 331m 75m 6904 S 20 0.9 0:26.50 qemu-system-arm The guest CPU was fully loaded, but the iothread could still run mostly independent on a second core. Without the patch we don't get beyond 32206 jan 20 0 330m 73m 7036 R 82 0.9 1:06.00 qemu-system-arm 32204 jan 20 0 330m 73m 7036 S 21 0.9 0:17.03 qemu-system-arm We don't benefit significantly, though, when the guest is not fully loading a host CPU. Signed-off-by: Jan Kiszka <jan.kiszka@siemens.com> Message-Id: <1439220437-23957-10-git-send-email-fred.konrad@greensocs.com> [FK: Rebase, fix qemu_devices_reset deadlock, rm address_space_* mutex] Signed-off-by: KONRAD Frederic <fred.konrad@greensocs.com> [EGC: fixed iothread lock for cpu-exec IRQ handling] Signed-off-by: Emilio G. Cota <cota@braap.org> [AJB: -smp single-threaded fix, clean commit msg, BQL fixes] Signed-off-by: Alex Bennée <alex.bennee@linaro.org> Reviewed-by: Richard Henderson <rth@twiddle.net> Reviewed-by: Pranith Kumar <bobby.prani@gmail.com> [PM: target-arm changes] Acked-by: Peter Maydell <peter.maydell@linaro.org>
2017-02-23 18:29:11 +00:00
qemu_mutex_lock_iothread();
locked = true;
}
r = memory_region_dispatch_read(mr, physaddr,
&val, size, iotlbentry->attrs);
if (r != MEMTX_OK) {
cpu_transaction_failed(cpu, physaddr, addr, size, MMU_DATA_LOAD,
mmu_idx, iotlbentry->attrs, r, retaddr);
}
tcg: drop global lock during TCG code execution This finally allows TCG to benefit from the iothread introduction: Drop the global mutex while running pure TCG CPU code. Reacquire the lock when entering MMIO or PIO emulation, or when leaving the TCG loop. We have to revert a few optimization for the current TCG threading model, namely kicking the TCG thread in qemu_mutex_lock_iothread and not kicking it in qemu_cpu_kick. We also need to disable RAM block reordering until we have a more efficient locking mechanism at hand. Still, a Linux x86 UP guest and my Musicpal ARM model boot fine here. These numbers demonstrate where we gain something: 20338 jan 20 0 331m 75m 6904 R 99 0.9 0:50.95 qemu-system-arm 20337 jan 20 0 331m 75m 6904 S 20 0.9 0:26.50 qemu-system-arm The guest CPU was fully loaded, but the iothread could still run mostly independent on a second core. Without the patch we don't get beyond 32206 jan 20 0 330m 73m 7036 R 82 0.9 1:06.00 qemu-system-arm 32204 jan 20 0 330m 73m 7036 S 21 0.9 0:17.03 qemu-system-arm We don't benefit significantly, though, when the guest is not fully loading a host CPU. Signed-off-by: Jan Kiszka <jan.kiszka@siemens.com> Message-Id: <1439220437-23957-10-git-send-email-fred.konrad@greensocs.com> [FK: Rebase, fix qemu_devices_reset deadlock, rm address_space_* mutex] Signed-off-by: KONRAD Frederic <fred.konrad@greensocs.com> [EGC: fixed iothread lock for cpu-exec IRQ handling] Signed-off-by: Emilio G. Cota <cota@braap.org> [AJB: -smp single-threaded fix, clean commit msg, BQL fixes] Signed-off-by: Alex Bennée <alex.bennee@linaro.org> Reviewed-by: Richard Henderson <rth@twiddle.net> Reviewed-by: Pranith Kumar <bobby.prani@gmail.com> [PM: target-arm changes] Acked-by: Peter Maydell <peter.maydell@linaro.org>
2017-02-23 18:29:11 +00:00
if (locked) {
qemu_mutex_unlock_iothread();
}
return val;
}
static void io_writex(CPUArchState *env, CPUIOTLBEntry *iotlbentry,
int mmu_idx,
uint64_t val, target_ulong addr,
uintptr_t retaddr, int size)
{
CPUState *cpu = ENV_GET_CPU(env);
hwaddr physaddr = iotlbentry->addr;
MemoryRegion *mr = iotlb_to_region(cpu, physaddr, iotlbentry->attrs);
tcg: drop global lock during TCG code execution This finally allows TCG to benefit from the iothread introduction: Drop the global mutex while running pure TCG CPU code. Reacquire the lock when entering MMIO or PIO emulation, or when leaving the TCG loop. We have to revert a few optimization for the current TCG threading model, namely kicking the TCG thread in qemu_mutex_lock_iothread and not kicking it in qemu_cpu_kick. We also need to disable RAM block reordering until we have a more efficient locking mechanism at hand. Still, a Linux x86 UP guest and my Musicpal ARM model boot fine here. These numbers demonstrate where we gain something: 20338 jan 20 0 331m 75m 6904 R 99 0.9 0:50.95 qemu-system-arm 20337 jan 20 0 331m 75m 6904 S 20 0.9 0:26.50 qemu-system-arm The guest CPU was fully loaded, but the iothread could still run mostly independent on a second core. Without the patch we don't get beyond 32206 jan 20 0 330m 73m 7036 R 82 0.9 1:06.00 qemu-system-arm 32204 jan 20 0 330m 73m 7036 S 21 0.9 0:17.03 qemu-system-arm We don't benefit significantly, though, when the guest is not fully loading a host CPU. Signed-off-by: Jan Kiszka <jan.kiszka@siemens.com> Message-Id: <1439220437-23957-10-git-send-email-fred.konrad@greensocs.com> [FK: Rebase, fix qemu_devices_reset deadlock, rm address_space_* mutex] Signed-off-by: KONRAD Frederic <fred.konrad@greensocs.com> [EGC: fixed iothread lock for cpu-exec IRQ handling] Signed-off-by: Emilio G. Cota <cota@braap.org> [AJB: -smp single-threaded fix, clean commit msg, BQL fixes] Signed-off-by: Alex Bennée <alex.bennee@linaro.org> Reviewed-by: Richard Henderson <rth@twiddle.net> Reviewed-by: Pranith Kumar <bobby.prani@gmail.com> [PM: target-arm changes] Acked-by: Peter Maydell <peter.maydell@linaro.org>
2017-02-23 18:29:11 +00:00
bool locked = false;
MemTxResult r;
physaddr = (physaddr & TARGET_PAGE_MASK) + addr;
if (mr != &io_mem_rom && mr != &io_mem_notdirty && !cpu->can_do_io) {
cpu_io_recompile(cpu, retaddr);
}
cpu->mem_io_vaddr = addr;
cpu->mem_io_pc = retaddr;
tcg: drop global lock during TCG code execution This finally allows TCG to benefit from the iothread introduction: Drop the global mutex while running pure TCG CPU code. Reacquire the lock when entering MMIO or PIO emulation, or when leaving the TCG loop. We have to revert a few optimization for the current TCG threading model, namely kicking the TCG thread in qemu_mutex_lock_iothread and not kicking it in qemu_cpu_kick. We also need to disable RAM block reordering until we have a more efficient locking mechanism at hand. Still, a Linux x86 UP guest and my Musicpal ARM model boot fine here. These numbers demonstrate where we gain something: 20338 jan 20 0 331m 75m 6904 R 99 0.9 0:50.95 qemu-system-arm 20337 jan 20 0 331m 75m 6904 S 20 0.9 0:26.50 qemu-system-arm The guest CPU was fully loaded, but the iothread could still run mostly independent on a second core. Without the patch we don't get beyond 32206 jan 20 0 330m 73m 7036 R 82 0.9 1:06.00 qemu-system-arm 32204 jan 20 0 330m 73m 7036 S 21 0.9 0:17.03 qemu-system-arm We don't benefit significantly, though, when the guest is not fully loading a host CPU. Signed-off-by: Jan Kiszka <jan.kiszka@siemens.com> Message-Id: <1439220437-23957-10-git-send-email-fred.konrad@greensocs.com> [FK: Rebase, fix qemu_devices_reset deadlock, rm address_space_* mutex] Signed-off-by: KONRAD Frederic <fred.konrad@greensocs.com> [EGC: fixed iothread lock for cpu-exec IRQ handling] Signed-off-by: Emilio G. Cota <cota@braap.org> [AJB: -smp single-threaded fix, clean commit msg, BQL fixes] Signed-off-by: Alex Bennée <alex.bennee@linaro.org> Reviewed-by: Richard Henderson <rth@twiddle.net> Reviewed-by: Pranith Kumar <bobby.prani@gmail.com> [PM: target-arm changes] Acked-by: Peter Maydell <peter.maydell@linaro.org>
2017-02-23 18:29:11 +00:00
if (mr->global_locking && !qemu_mutex_iothread_locked()) {
tcg: drop global lock during TCG code execution This finally allows TCG to benefit from the iothread introduction: Drop the global mutex while running pure TCG CPU code. Reacquire the lock when entering MMIO or PIO emulation, or when leaving the TCG loop. We have to revert a few optimization for the current TCG threading model, namely kicking the TCG thread in qemu_mutex_lock_iothread and not kicking it in qemu_cpu_kick. We also need to disable RAM block reordering until we have a more efficient locking mechanism at hand. Still, a Linux x86 UP guest and my Musicpal ARM model boot fine here. These numbers demonstrate where we gain something: 20338 jan 20 0 331m 75m 6904 R 99 0.9 0:50.95 qemu-system-arm 20337 jan 20 0 331m 75m 6904 S 20 0.9 0:26.50 qemu-system-arm The guest CPU was fully loaded, but the iothread could still run mostly independent on a second core. Without the patch we don't get beyond 32206 jan 20 0 330m 73m 7036 R 82 0.9 1:06.00 qemu-system-arm 32204 jan 20 0 330m 73m 7036 S 21 0.9 0:17.03 qemu-system-arm We don't benefit significantly, though, when the guest is not fully loading a host CPU. Signed-off-by: Jan Kiszka <jan.kiszka@siemens.com> Message-Id: <1439220437-23957-10-git-send-email-fred.konrad@greensocs.com> [FK: Rebase, fix qemu_devices_reset deadlock, rm address_space_* mutex] Signed-off-by: KONRAD Frederic <fred.konrad@greensocs.com> [EGC: fixed iothread lock for cpu-exec IRQ handling] Signed-off-by: Emilio G. Cota <cota@braap.org> [AJB: -smp single-threaded fix, clean commit msg, BQL fixes] Signed-off-by: Alex Bennée <alex.bennee@linaro.org> Reviewed-by: Richard Henderson <rth@twiddle.net> Reviewed-by: Pranith Kumar <bobby.prani@gmail.com> [PM: target-arm changes] Acked-by: Peter Maydell <peter.maydell@linaro.org>
2017-02-23 18:29:11 +00:00
qemu_mutex_lock_iothread();
locked = true;
}
r = memory_region_dispatch_write(mr, physaddr,
val, size, iotlbentry->attrs);
if (r != MEMTX_OK) {
cpu_transaction_failed(cpu, physaddr, addr, size, MMU_DATA_STORE,
mmu_idx, iotlbentry->attrs, r, retaddr);
}
tcg: drop global lock during TCG code execution This finally allows TCG to benefit from the iothread introduction: Drop the global mutex while running pure TCG CPU code. Reacquire the lock when entering MMIO or PIO emulation, or when leaving the TCG loop. We have to revert a few optimization for the current TCG threading model, namely kicking the TCG thread in qemu_mutex_lock_iothread and not kicking it in qemu_cpu_kick. We also need to disable RAM block reordering until we have a more efficient locking mechanism at hand. Still, a Linux x86 UP guest and my Musicpal ARM model boot fine here. These numbers demonstrate where we gain something: 20338 jan 20 0 331m 75m 6904 R 99 0.9 0:50.95 qemu-system-arm 20337 jan 20 0 331m 75m 6904 S 20 0.9 0:26.50 qemu-system-arm The guest CPU was fully loaded, but the iothread could still run mostly independent on a second core. Without the patch we don't get beyond 32206 jan 20 0 330m 73m 7036 R 82 0.9 1:06.00 qemu-system-arm 32204 jan 20 0 330m 73m 7036 S 21 0.9 0:17.03 qemu-system-arm We don't benefit significantly, though, when the guest is not fully loading a host CPU. Signed-off-by: Jan Kiszka <jan.kiszka@siemens.com> Message-Id: <1439220437-23957-10-git-send-email-fred.konrad@greensocs.com> [FK: Rebase, fix qemu_devices_reset deadlock, rm address_space_* mutex] Signed-off-by: KONRAD Frederic <fred.konrad@greensocs.com> [EGC: fixed iothread lock for cpu-exec IRQ handling] Signed-off-by: Emilio G. Cota <cota@braap.org> [AJB: -smp single-threaded fix, clean commit msg, BQL fixes] Signed-off-by: Alex Bennée <alex.bennee@linaro.org> Reviewed-by: Richard Henderson <rth@twiddle.net> Reviewed-by: Pranith Kumar <bobby.prani@gmail.com> [PM: target-arm changes] Acked-by: Peter Maydell <peter.maydell@linaro.org>
2017-02-23 18:29:11 +00:00
if (locked) {
qemu_mutex_unlock_iothread();
}
}
/* Return true if ADDR is present in the victim tlb, and has been copied
back to the main tlb. */
static bool victim_tlb_hit(CPUArchState *env, size_t mmu_idx, size_t index,
size_t elt_ofs, target_ulong page)
{
size_t vidx;
for (vidx = 0; vidx < CPU_VTLB_SIZE; ++vidx) {
CPUTLBEntry *vtlb = &env->tlb_v_table[mmu_idx][vidx];
target_ulong cmp = *(target_ulong *)((uintptr_t)vtlb + elt_ofs);
if (cmp == page) {
/* Found entry in victim tlb, swap tlb and iotlb. */
CPUTLBEntry tmptlb, *tlb = &env->tlb_table[mmu_idx][index];
copy_tlb_helper(&tmptlb, tlb, false);
copy_tlb_helper(tlb, vtlb, true);
copy_tlb_helper(vtlb, &tmptlb, true);
CPUIOTLBEntry tmpio, *io = &env->iotlb[mmu_idx][index];
CPUIOTLBEntry *vio = &env->iotlb_v[mmu_idx][vidx];
tmpio = *io; *io = *vio; *vio = tmpio;
return true;
}
}
return false;
}
/* Macro to call the above, with local variables from the use context. */
#define VICTIM_TLB_HIT(TY, ADDR) \
victim_tlb_hit(env, mmu_idx, index, offsetof(CPUTLBEntry, TY), \
(ADDR) & TARGET_PAGE_MASK)
/* NOTE: this function can trigger an exception */
/* NOTE2: the returned address is not exactly the physical address: it
* is actually a ram_addr_t (in system mode; the user mode emulation
* version of this function returns a guest virtual address).
*/
tb_page_addr_t get_page_addr_code(CPUArchState *env, target_ulong addr)
{
int mmu_idx, index, pd;
void *p;
MemoryRegion *mr;
CPUState *cpu = ENV_GET_CPU(env);
CPUIOTLBEntry *iotlbentry;
hwaddr physaddr;
index = (addr >> TARGET_PAGE_BITS) & (CPU_TLB_SIZE - 1);
mmu_idx = cpu_mmu_index(env, true);
if (unlikely(env->tlb_table[mmu_idx][index].addr_code !=
(addr & (TARGET_PAGE_MASK | TLB_INVALID_MASK)))) {
if (!VICTIM_TLB_HIT(addr_read, addr)) {
tlb_fill(ENV_GET_CPU(env), addr, MMU_INST_FETCH, mmu_idx, 0);
}
}
iotlbentry = &env->iotlb[mmu_idx][index];
pd = iotlbentry->addr & ~TARGET_PAGE_MASK;
mr = iotlb_to_region(cpu, pd, iotlbentry->attrs);
if (memory_region_is_unassigned(mr)) {
qemu_mutex_lock_iothread();
if (memory_region_request_mmio_ptr(mr, addr)) {
qemu_mutex_unlock_iothread();
/* A MemoryRegion is potentially added so re-run the
* get_page_addr_code.
*/
return get_page_addr_code(env, addr);
}
qemu_mutex_unlock_iothread();
/* Give the new-style cpu_transaction_failed() hook first chance
* to handle this.
* This is not the ideal place to detect and generate CPU
* exceptions for instruction fetch failure (for instance
* we don't know the length of the access that the CPU would
* use, and it would be better to go ahead and try the access
* and use the MemTXResult it produced). However it is the
* simplest place we have currently available for the check.
*/
physaddr = (iotlbentry->addr & TARGET_PAGE_MASK) + addr;
cpu_transaction_failed(cpu, physaddr, addr, 0, MMU_INST_FETCH, mmu_idx,
iotlbentry->attrs, MEMTX_DECODE_ERROR, 0);
cpu_unassigned_access(cpu, addr, false, true, 0, 4);
/* The CPU's unassigned access hook might have longjumped out
* with an exception. If it didn't (or there was no hook) then
* we can't proceed further.
*/
report_bad_exec(cpu, addr);
exit(1);
}
p = (void *)((uintptr_t)addr + env->tlb_table[mmu_idx][index].addend);
return qemu_ram_addr_from_host_nofail(p);
}
/* Probe for whether the specified guest write access is permitted.
* If it is not permitted then an exception will be taken in the same
* way as if this were a real write access (and we will not return).
* Otherwise the function will return, and there will be a valid
* entry in the TLB for this access.
*/
void probe_write(CPUArchState *env, target_ulong addr, int mmu_idx,
uintptr_t retaddr)
{
int index = (addr >> TARGET_PAGE_BITS) & (CPU_TLB_SIZE - 1);
target_ulong tlb_addr = env->tlb_table[mmu_idx][index].addr_write;
if ((addr & TARGET_PAGE_MASK)
!= (tlb_addr & (TARGET_PAGE_MASK | TLB_INVALID_MASK))) {
/* TLB entry is for a different page */
if (!VICTIM_TLB_HIT(addr_write, addr)) {
tlb_fill(ENV_GET_CPU(env), addr, MMU_DATA_STORE, mmu_idx, retaddr);
}
}
}
/* Probe for a read-modify-write atomic operation. Do not allow unaligned
* operations, or io operations to proceed. Return the host address. */
static void *atomic_mmu_lookup(CPUArchState *env, target_ulong addr,
TCGMemOpIdx oi, uintptr_t retaddr)
{
size_t mmu_idx = get_mmuidx(oi);
size_t index = (addr >> TARGET_PAGE_BITS) & (CPU_TLB_SIZE - 1);
CPUTLBEntry *tlbe = &env->tlb_table[mmu_idx][index];
target_ulong tlb_addr = tlbe->addr_write;
TCGMemOp mop = get_memop(oi);
int a_bits = get_alignment_bits(mop);
int s_bits = mop & MO_SIZE;
/* Adjust the given return address. */
retaddr -= GETPC_ADJ;
/* Enforce guest required alignment. */
if (unlikely(a_bits > 0 && (addr & ((1 << a_bits) - 1)))) {
/* ??? Maybe indicate atomic op to cpu_unaligned_access */
cpu_unaligned_access(ENV_GET_CPU(env), addr, MMU_DATA_STORE,
mmu_idx, retaddr);
}
/* Enforce qemu required alignment. */
if (unlikely(addr & ((1 << s_bits) - 1))) {
/* We get here if guest alignment was not requested,
or was not enforced by cpu_unaligned_access above.
We might widen the access and emulate, but for now
mark an exception and exit the cpu loop. */
goto stop_the_world;
}
/* Check TLB entry and enforce page permissions. */
if ((addr & TARGET_PAGE_MASK)
!= (tlb_addr & (TARGET_PAGE_MASK | TLB_INVALID_MASK))) {
if (!VICTIM_TLB_HIT(addr_write, addr)) {
tlb_fill(ENV_GET_CPU(env), addr, MMU_DATA_STORE, mmu_idx, retaddr);
}
tlb_addr = tlbe->addr_write;
}
/* Check notdirty */
if (unlikely(tlb_addr & TLB_NOTDIRTY)) {
tlb_set_dirty(ENV_GET_CPU(env), addr);
tlb_addr = tlb_addr & ~TLB_NOTDIRTY;
}
/* Notice an IO access */
if (unlikely(tlb_addr & ~TARGET_PAGE_MASK)) {
/* There's really nothing that can be done to
support this apart from stop-the-world. */
goto stop_the_world;
}
/* Let the guest notice RMW on a write-only page. */
if (unlikely(tlbe->addr_read != tlb_addr)) {
tlb_fill(ENV_GET_CPU(env), addr, MMU_DATA_LOAD, mmu_idx, retaddr);
/* Since we don't support reads and writes to different addresses,
and we do have the proper page loaded for write, this shouldn't
ever return. But just in case, handle via stop-the-world. */
goto stop_the_world;
}
return (void *)((uintptr_t)addr + tlbe->addend);
stop_the_world:
cpu_loop_exit_atomic(ENV_GET_CPU(env), retaddr);
}
#ifdef TARGET_WORDS_BIGENDIAN
# define TGT_BE(X) (X)
# define TGT_LE(X) BSWAP(X)
#else
# define TGT_BE(X) BSWAP(X)
# define TGT_LE(X) (X)
#endif
#define MMUSUFFIX _mmu
#define DATA_SIZE 1
#include "softmmu_template.h"
#define DATA_SIZE 2
#include "softmmu_template.h"
#define DATA_SIZE 4
#include "softmmu_template.h"
#define DATA_SIZE 8
#include "softmmu_template.h"
/* First set of helpers allows passing in of OI and RETADDR. This makes
them callable from other helpers. */
#define EXTRA_ARGS , TCGMemOpIdx oi, uintptr_t retaddr
#define ATOMIC_NAME(X) \
HELPER(glue(glue(glue(atomic_ ## X, SUFFIX), END), _mmu))
#define ATOMIC_MMU_LOOKUP atomic_mmu_lookup(env, addr, oi, retaddr)
#define DATA_SIZE 1
#include "atomic_template.h"
#define DATA_SIZE 2
#include "atomic_template.h"
#define DATA_SIZE 4
#include "atomic_template.h"
#ifdef CONFIG_ATOMIC64
#define DATA_SIZE 8
#include "atomic_template.h"
#endif
#ifdef CONFIG_ATOMIC128
#define DATA_SIZE 16
#include "atomic_template.h"
#endif
/* Second set of helpers are directly callable from TCG as helpers. */
#undef EXTRA_ARGS
#undef ATOMIC_NAME
#undef ATOMIC_MMU_LOOKUP
#define EXTRA_ARGS , TCGMemOpIdx oi
#define ATOMIC_NAME(X) HELPER(glue(glue(atomic_ ## X, SUFFIX), END))
#define ATOMIC_MMU_LOOKUP atomic_mmu_lookup(env, addr, oi, GETPC())
#define DATA_SIZE 1
#include "atomic_template.h"
#define DATA_SIZE 2
#include "atomic_template.h"
#define DATA_SIZE 4
#include "atomic_template.h"
#ifdef CONFIG_ATOMIC64
#define DATA_SIZE 8
#include "atomic_template.h"
#endif
/* Code access functions. */
#undef MMUSUFFIX
#define MMUSUFFIX _cmmu
#undef GETPC
#define GETPC() ((uintptr_t)0)
#define SOFTMMU_CODE_ACCESS
#define DATA_SIZE 1
#include "softmmu_template.h"
#define DATA_SIZE 2
#include "softmmu_template.h"
#define DATA_SIZE 4
#include "softmmu_template.h"
#define DATA_SIZE 8
#include "softmmu_template.h"