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storesets paper.
This patch improves performance by as much as 10% on some spec benchmarks.
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SEV instructions were originally implemented to cause asynchronous squashes
via the generateTCSquash() function in the O3 pipeline when updating the
SEV_MAILBOX miscReg. This caused race conditions between CPUs in an MP system
that would lead to a pipeline either going inactive indefinitely or not being
able to commit squashed instructions. Fixed SEV instructions to behave like
interrupts and cause synchronous sqaushes inside the pipeline, eliminating
the race conditions. Also fixed up the semantics of the WFE instruction to
behave as documented in the ARMv7 ISA description to not sleep if SEV_MAILBOX=1
or unmasked interrupts are pending.
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Two issues are fixed in this patch:
1. The load and store pc passed to the predictor are passed in reverse order.
2. The flag indicating that a barrier is inflight was never cleared when
the barrier was squashed instead of committed. This made all load insts
dependent on a non-existent barrier in-flight.
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Change the way instructions are squashed on memory ordering violations
to squash the violator and younger instructions, not all instructions
that are younger than the instruction they violated (no reason to throw
away valid work).
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It's possible (though until now very unlikely) for fetchAddr to get out of
sync with the actual PC of the current instruction. This change forcefull
resets fetchAddr at the end of every instruction.
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Until now, the only reason a macroop would be left was because it ended at a
microop marked as the last microop. In O3 with branch prediction, it's
possible for the branch predictor to have entries which originally came from
different instructions which happened to have the same RIP. This could
theoretically happen in many ways, but it was encountered specifically when
different programs in different address spaces ran one after the other in
X86_FS.
What would happen in that case was that the macroop would continue to be
looped over and microops fetched from it until it reached the last microop
even though the macropc had moved out from under it. If things lined up
properly, this could mean that the end bytes of an instruction actually fell
into the instruction sized block of memory after the one in the predecoder.
The fetch loop implicitly assumes that the last instruction sized chunk of
memory processed was the last one needed for the instruction it just finished
executing. It would then tell the predecoder to move to an offset within the
bytes it was given that is larger than those bytes, and that would trip an
assert in the x86 predecoder.
This change fixes this problem by making fetch stop processing the current
macroop if the address it should be fetching from changed when the PC is
updated. That happens when the last microop was reached because the instruction
handled it properly, and it also catches the case where the branch predictor
makes fetch do a macro level branch when it shouldn't.
The check of isLastMicroop is retained because otherwise, a macroop that
branches back to itself would act like a single, long macroop instead of
multiple instances of the same microop. There may be situations (which may
turn out to be purely hypothetical) where that matters.
This also fixes a relatively minor issue where the curMacroop variable would
be set to NULL immediately after seeing that a microop was the last one before
curMacroop was used to build the dyninst. The traceData structure would have a
NULL pointer to the macroop for that microop.
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Before this change, the commit stage would wait until the ROB and store queue
were empty before recognizing an interrupt. The fetch stage would stop
generating instructions at an appropriate point, so commit would then wait
until a valid time to interrupt the instruction stream. Instructions might be
in flight after fetch but not the in the ROB or store queue (in rename, for
instance), so this change makes commit wait until all in flight instructions
are finished.
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Let squahsed and deferred instructions issue so they don't accumulate and clog
up the CPU.
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This constructor assumes that the ExtMachInst can be decoded directly into a
StaticInst that's useful to execute. With the advent of microcoded
instructions that's no longer true.
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When fetching from the microcode ROM, if the PC is set so that it isn't in the
cache block that's been fetched the CPU will get stuck. The fetch stage
notices that it's in the ROM so it doesn't try to fetch from the current PC.
It then later notices that it's outside of the current cache block so it skips
generating instructions expecting to continue once the right bytes have been
fetched. This change lets the fetch stage attempt to generate instructions,
and only checks if the bytes it's going to use are valid if it's really going
to use them.
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Implemented a pipeline activity viewer as a python script (util/o3-pipeview.py)
and modified O3 code base to support an extra trace flag (O3PipeView) for
generating traces to be used as inputs by the tool.
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Fixed up the patch from Yasuko Watanabe that enabled pipelining of fetch accessess to
icache to work with recent changes to main repository.
Also added in ability for fetch stage to delay issuing the fault carrying
nop when a pipeline fetch causes a fault and no fetch bandwidth is available
until the next cycle.
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Calculation of offset to copy from storeQueue[idx].data structure for load to
store forwarding fixed to be difference in bytes between store and load virtual
addresses. Previous method would induce bug where a load would index into
buffer at the wrong location.
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If a split load fails on a blocked cache wbOutstanding can be decremented
twice if the first part of the split load succeeds and the second part fails.
Condition the decrementing on not having completed the first part of the load.
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This patch fixes two problems with the O3 cpu model. The first is an issue
with an instruction fetch causing a fault on the next address while the
current macro-op is being issued. This happens when the micro-ops exceed
the fetch bandwdith and then on the next cycle the fetch stage attempts
to issue a request to the next line while it still has micro-ops to issue
if the next line faults a fault is attached to a micro-op in the currently
executing macro-op rather than a "nop" from the next instruction block.
This leads to an instruction incorrectly faulting when on fetch when
it had no reason to fault.
A similar problem occurs with interrupts. When an interrupt occurs the
fetch stage nominally stops issuing instructions immediately. This is incorrect
in the case of a macro-op as the current location might not be interruptable.
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Instructions that load an address and are control instructions can
execute down the wrong path if they were predicted correctly and then
instructions following them are squashed. If an instruction is a
memory and control op use the predicted address for the next PC instead
of just advancing the PC. Without this change NPC is used for the next
instruction, but predPC is used to verify that the branch was successful
so the wrong path is silently executed.
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If an nonspeculative instruction has a fault it might not be in the
nonSpecInsts map.
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At the same time, rename the trace flags to debug flags since they
have broader usage than simply tracing. This means that
--trace-flags is now --debug-flags and --trace-help is now --debug-help
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This change fixes a small bug in the arm copyRegs() code where some registers
wouldn't be copied if the processor was in a mode other than MODE_USER.
Additionally, this change simplifies the way the O3 switchCpu code works by
utilizing TheISA::copyRegs() to copy the required context information
rather than the adhoc copying that goes on in the CPU model. The current code
makes assumptions about the visibility of int and float registers that aren't
true for all architectures in FS mode.
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Consolidate all code to handle ITSTATE in the PCState object rather than
touching a variety of structures/objects.
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The comment in the code suggests that the checking granularity should be 16
bytes, however in reality the shift by 8 is 256 bytes which seems much
larger than required.
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Get rid of unused members and use base types rather than derrived values
where possible to limit amount of state.
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non-cache.
This change fixes the problem for all the cases we actively use. If you want to try
more creative I/O device attachments (E.g. sharing an L2), this won't work. You
would need another level of caching between the I/O device and the cache
(which you actually need anyway with our current code to make sure writes
propagate). This is required so that you can mark the cache in between as
top level and it won't try to send ownership of a block to the I/O device.
Asserts have been added that should catch any issues.
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Without this change the a store can be issued to the cache multiple times.
If this case occurs when the l1 cache is out of mshrs (and thus blocked)
the processor will never make forward progress because each cycle it will
send a single request using the recently freed mshr and not completing the
multipart store. This will continue forever.
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Fixes two of the SMT fetch policies in O3CPU that were returning the count
of instructions in the IQ or LSQ rather than the thread ID to fetch from.
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If there is an outstanding table walk and no other activity in the CPU
it can go to sleep and never wake up. This change makes the instruction
queue always active if the CPU is waiting for a store to translate.
If Gabe changes the way this code works then the below should be removed
as indicated by the todo.
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In this case we need to throw away the TLB miss, not assume it was the
one we were waiting for.
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When a table walk is initiated by the fetch stage, the CPU can
potentially move to the idle state and never wake up.
The fetch stage must call cpu->wakeCPU() when a translation completes
(in finishTranslation()).
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Some ISAs (like ARM) relies on hardware page table walkers. For those ISAs,
when a TLB miss occurs, initiateTranslation() can return with NoFault but with
the translation unfinished.
Instructions experiencing a delayed translation due to a hardware page table
walk are deferred until the translation completes and kept into the IQ. In
order to keep track of them, the IQ has been augmented with a queue of the
outstanding delayed memory instructions. When their translation completes,
instructions are re-executed (only their initiateAccess() was already
executed; their DTB translation is now skipped). The IEW stage has been
modified to support such a 2-pass execution.
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