This patch adds assertions that enforce that only invalidating snoops
will ever reach into the logic that tracks in-order load completion and
also invalidation of LL/SC (and MONITOR / MWAIT) monitors. Also adds
some comments to MSHR::replaceUpgrades().
Writes to locked memory addresses (LLSC) did not wake up the locking
CPU. This can lead to deadlocks on multi-core runs. In AtomicSimpleCPU,
recvAtomicSnoop was checking if the incoming packet was an invalidation
(isInvalidate) and only then handled a locked snoop. But, writes are
seen instead of invalidates when running without caches (fast-forward
configurations). As as simple fix, now handleLockedSnoop is also called
even if the incoming snoop packet are from writes.
This patch changes how the cache determines if snoops should be
forwarded from the memory side to the CPU side. Instead of having a
parameter, the cache now looks at the port connected on the CPU side,
and if it is a snooping port, then snoops are forwarded. Less error
prone, and less parameters to worry about.
The patch also tidies up the CPU classes to ensure that their I-side
port is not snooping by removing overrides to the snoop request
handler, such that snoop requests will panic via the default
MasterPort implement
The read() function merely initiates a memory read operation; the
data doesn't arrive until the access completes and a response packet
is received from the memory system. Thus there's no need to provide
a data pointer; its existence is historical.
Getting this pointer out of this internal o3 interface sets the
stage for similar cleanup in the ExecContext interface. Also
found that we were pointlessly setting the contents at this pointer
on a store forward (the useful memcpy happens just a few lines
below the deleted one).
This patch changes the name of a bunch of packet flags and MSHR member
functions and variables to make the coherency protocol easier to
understand. In addition the patch adds and updates lots of
descriptions, explicitly spelling out assumptions.
The following name changes are made:
* the packet memInhibit flag is renamed to cacheResponding
* the packet sharedAsserted flag is renamed to hasSharers
* the packet NeedsExclusive attribute is renamed to NeedsWritable
* the packet isSupplyExclusive is renamed responderHadWritable
* the MSHR pendingDirty is renamed to pendingModified
The cache states, Modified, Owned, Exclusive, Shared are also called
out in the cache and MSHR code to make it easier to understand.
This patch adds support to optionally capture the virtual address and asid
for load/store instructions in the elastic traces. If they are present in
the traces, Trace CPU will set those fields of the request during replay.
This patch replaces the booleans that specified the elastic trace record
type with an enum type. The source of change is the proto message for
elastic trace where the enum is introduced. The struct definitions in the
elastic trace probe listener as well as the Trace CPU replace the boleans
with the proto message enum.
The patch does not impact functionality, but traces are not compatible with
previous version. This is preparation for adding new types of records in
subsequent patches.
The elastic trace is a type of probe listener and listens to probe points
in multiple stages of the O3CPU. The notify method is called on a probe
point typically when an instruction successfully progresses through that
stage.
As different listener methods mapped to the different probe points execute,
relevant information about the instruction, e.g. timestamps and register
accesses, are captured and stored in temporary InstExecInfo class objects.
When the instruction progresses through the commit stage, the timing and the
dependency information about the instruction is finalised and encapsulated in
a struct called TraceInfo. TraceInfo objects are collected in a list instead
of writing them out to the trace file one a time. This is required as the
trace is processed in chunks to evaluate order dependencies and computational
delay in case an instruction does not have any register dependencies. By this
we achieve a simpler algorithm during replay because every record in the
trace can be hooked onto a record in its past. The instruction dependency
trace is written out as a protobuf format file. A second trace containing
fetch requests at absolute timestamps is written to a separate protobuf
format file.
If the instruction is not executed then it is not added to the trace.
The code checks if the instruction had a fault, if it predicated
false and thus previous register values were restored or if it was a
load/store that did not have a request (e.g. when the size of the
request is zero). In all these cases the instruction is set as
executed by the Execute stage and is picked up by the commit probe
listener. But a request is not issued and registers are not written.
So practically, skipping these should not hurt the dependency modelling.
If squashing results in squashing younger instructions, it may happen that
the squash probe discards the inst and removes it from the temporary
store but execute stage deals with the instruction in the next cycle which
results in the execute probe seeing this inst as 'new' inst. A sequence
number of the last processed trace record is used to trap these cases and
not add to the temporary store.
The elastic instruction trace and fetch request trace can be read in and
played back by the TraceCPU.
This patch adds probe points in Fetch, IEW, Rename and Commit stages as follows.
A probe point is added in the Fetch stage for probing when a fetch request is
sent. Notify is fired on the probe point when a request is sent succesfully in
the first attempt as well as on a retry attempt.
Probe points are added in the IEW stage when an instruction begins to execute
and when execution is complete. This points can be used for monitoring the
execution time of an instruction.
Probe points are added in the Rename stage to probe renaming of source and
destination registers and when there is squashing. These probe points can be
used to track register dependencies and remove when there is squashing.
A probe point for squashing is added in Commit to probe squashed instructions.
The assert in lsq_unit_impl.hh line 963 needs pktPending to be initialized to
NULL (I got the assertion failure several times without the fix).
Committed by: Nilay Vaish <nilay@cs.wisc.edu>
This patch adds explicit overrides as this is now required when using
"-Wall" with clang >= 3.5, the latter now part of the most recent
XCode. The patch consequently removes "virtual" for those methods
where "override" is added. The latter should be enough of an
indication.
As part of this patch, a few minor issues that clang >= 3.5 complains
about are also resolved (unused methods and variables).
This patch moves away from using M5_ATTR_OVERRIDE and the m5::hashmap
(and similar) abstractions, as these are no longer needed with gcc 4.7
and clang 3.1 as minimum compiler versions.
The decoder is responsible for splitting instructions in micro
operations (uops). Given that different micro architectures may split
operations differently, this patch allows to specify which micro
architecture each isa implements, so different cores in the system can
split instructions differently, also decoupling uop splitting
(microArch) from ISA (Arch). This is done making the decodification
calls templates that receive a type 'DecoderFlavour' that maps the
name of the operation to the class that implements it. This way there
is only one selection point (converting the command line enum to the
appropriate DecodeFeatures object). In addition, there is no explicit
code replication: template instantiation hides that, and the compiler
should be able to resolve a number of things at compile-time.
This patch enables instructions in LSQ to track two physical addresses for
corresponding two split requests. Later, the information is used in
checksnoop() to search for/invalidate the corresponding LD instructions.
The current implementation has kept track of only the physical address that is
referenced by the first split request. Thus, for checksnoop(), the line
accessed by the second request has not been considered, causing potential
correctness issues.
Committed by: Nilay Vaish <nilay@cs.wisc.edu>
This adds a vector register type. The type is defined as a std::array of a
fixed number of uint64_ts. The isa_parser.py has been modified to parse vector
register operands and generate the required code. Different cpus have vector
register files now.
The drain() call currently passes around a DrainManager pointer, which
is now completely pointless since there is only ever one global
DrainManager in the system. It also contains vestiges from the time
when SimObjects had to keep track of their child objects that needed
draining.
This changeset moves all of the DrainState handling to the Drainable
base class and changes the drain() and drainResume() calls to reflect
this. Particularly, the drain() call has been updated to take no
parameters (the DrainManager argument isn't needed) and return a
DrainState instead of an unsigned integer (there is no point returning
anything other than 0 or 1 any more). Drainable objects should return
either DrainState::Draining (equivalent to returning 1 in the old
system) if they need more time to drain or DrainState::Drained
(equivalent to returning 0 in the old system) if they are already in a
consistent state. Returning DrainState::Running is considered an
error.
Drain done signalling is now done through the signalDrainDone() method
in the Drainable class instead of using the DrainManager directly. The
new call checks if the state of the object is DrainState::Draining
before notifying the drain manager. This means that it is safe to call
signalDrainDone() without first checking if the simulator has
requested draining. The intention here is to reduce the code needed to
implement draining in simple objects.
The drain state enum is currently a part of the Drainable
interface. The same state machine will be used by the DrainManager to
identify the global state of the simulator. Make the drain state a
global typed enum to better cater for this usage scenario.
Objects that are can be serialized are supposed to inherit from the
Serializable class. This class is meant to provide a unified API for
such objects. However, so far it has mainly been used by SimObjects
due to some fundamental design limitations. This changeset redesigns
to the serialization interface to make it more generic and hide the
underlying checkpoint storage. Specifically:
* Add a set of APIs to serialize into a subsection of the current
object. Previously, objects that needed this functionality would
use ad-hoc solutions using nameOut() and section name
generation. In the new world, an object that implements the
interface has the methods serializeSection() and
unserializeSection() that serialize into a named /subsection/ of
the current object. Calling serialize() serializes an object into
the current section.
* Move the name() method from Serializable to SimObject as it is no
longer needed for serialization. The fully qualified section name
is generated by the main serialization code on the fly as objects
serialize sub-objects.
* Add a scoped ScopedCheckpointSection helper class. Some objects
need to serialize data structures, that are not deriving from
Serializable, into subsections. Previously, this was done using
nameOut() and manual section name generation. To simplify this,
this changeset introduces a ScopedCheckpointSection() helper
class. When this class is instantiated, it adds a new /subsection/
and subsequent serialization calls during the lifetime of this
helper class happen inside this section (or a subsection in case
of nested sections).
* The serialize() call is now const which prevents accidental state
manipulation during serialization. Objects that rely on modifying
state can use the serializeOld() call instead. The default
implementation simply calls serialize(). Note: The old-style calls
need to be explicitly called using the
serializeOld()/serializeSectionOld() style APIs. These are used by
default when serializing SimObjects.
* Both the input and output checkpoints now use their own named
types. This hides underlying checkpoint implementation from
objects that need checkpointing and makes it easier to change the
underlying checkpoint storage code.
Three minor issues are resolved:
1. Apparently gcc 5.1 does not like negation of booleans followed by
bitwise AND.
2. Somehow the compiler also gets confused and warns about
NoopMachInst being unused (removing it causes compilation errors
though). Most likely a compiler bug.
3. There seems to be a number of instances where loop unrolling causes
false positives for the array-bounds check. For now, switch to
std::array. Potentially we could disable the warning for newer gcc
versions, but switching to std::array is probably a good move in
any case.
The Request::UNCACHEABLE flag currently has two different
functions. The first, and obvious, function is to prevent the memory
system from caching data in the request. The second function is to
prevent reordering and speculation in CPU models.
This changeset gives the order/speculation requirement a separate flag
(Request::STRICT_ORDER). This flag prevents CPU models from doing the
following optimizations:
* Speculation: CPU models are not allowed to issue speculative
loads.
* Write combining: CPU models and caches are not allowed to merge
writes to the same cache line.
Note: The memory system may still reorder accesses unless the
UNCACHEABLE flag is set. It is therefore expected that the
STRICT_ORDER flag is combined with the UNCACHEABLE flag to prevent
this behavior.
This patch takes a last step in fixing issues related to uncacheable
accesses. We do not separate uncacheable memory from uncacheable
devices, and in cases where it is really memory, there are valid
scenarios where we need to snoop since we do not support cache
maintenance instructions (yet). On snooping an uncacheable access we
thus provide data if possible. In essence this makes uncacheable
accesses IO coherent.
The snoop filter is also queried to steer the snoops, but not updated
since the uncacheable accesses do not allocate a block.
This patch fixes a recent issue with gcc 4.9 (and possibly more) being
convinced that indices outside the array bounds are used when
initialising the FUPool members.
Currently, each op class has a parameter issueLat that denotes the cycles after
which another op of the same class can be issued. As of now, this latency can
either be one cycle (fully pipelined) or same as execution latency of the op
(not at all pipelined). The fact that issueLat is a parameter of type Cycles
makes one believe that it can be set to any value. To avoid the confusion, the
parameter is being renamed as 'pipelined' with type boolean. If set to true,
the op would execute in a fully pipelined fashion. Otherwise, it would execute
in an unpipelined fashion.
This patch sets the default latency of the division microop to a single cycle
on x86. This is because the division instructions DIV and IDIV have been
implemented as loops of div microops, where each microop computes a single bit
of the quotient.
The o3 cpu instruction queue model uses the count variable to track the number
of unissued instructions in the queue. Previously, the squash method used
this variable to avoid executing the doSquash method when there were no
unissued instructions in the pipeline. A corner case problem exists when
only issued instructions exist in the pipeline and a squash occurs; the
doSquash code is not invoked and subsequently does not clean up state properly.
The totalInstructions counter is only incremented when the whole instruction is
commited and not on every microop. It was incorrectly reset in atomic and
timing cpus.
Committed by: Nilay Vaish <nilay@cs.wisc.edu>"
The variable is used in only one place and a whole new function setNextStatus()
has been defined just to compute the value of the variable. Instead of calling
the function, the value is now computed in the loop that preceded the function
call.
Now, prior to the renaming, the instruction requests the exact amount of
registers it will need, and the rename_map decides whether the instruction is
allowed to proceed or not.
This patch fixes a long-standing isue with the port flow
control. Before this patch the retry mechanism was shared between all
different packet classes. As a result, a snoop response could get
stuck behind a request waiting for a retry, even if the send/recv
functions were split. This caused message-dependent deadlocks in
stress-test scenarios.
The patch splits the retry into one per packet (message) class. Thus,
sendTimingReq has a corresponding recvReqRetry, sendTimingResp has
recvRespRetry etc. Most of the changes to the code involve simply
clarifying what type of request a specific object was accepting.
The biggest change in functionality is in the cache downstream packet
queue, facing the memory. This queue was shared by requests and snoop
responses, and it is now split into two queues, each with their own
flow control, but the same physical MasterPort. These changes fixes
the previously seen deadlocks.
