This implements most of the architected PMU events. The ones missing
are mostly the ones that depend on which level of the cache hierarchy
data is fetched from. The events implemented here, and their raw
event codes, are:
Floating-point operation completed (100f4)
Load completed (100fc)
Store completed (200f0)
Icache miss (200fc)
ITLB miss (100f6)
ITLB miss resolved (400fc)
Dcache load miss (400f0)
Dcache load miss resolved (300f8)
Dcache store miss (300f0)
DTLB miss (300fc)
DTLB miss resolved (200f6)
No instruction available and none being executed (100f8)
Instruction dispatched (200f2, 300f2, 400f2)
Taken branch instruction completed (200fa)
Branch mispredicted (400f6)
External interrupt taken (200f8)
Signed-off-by: Paul Mackerras <paulus@ozlabs.org>
At present the logic prevents any interrupts from being handled while
there is a load/store instruction (one that has unit=LDST) being
executed. However, load/store instructions can still get sent to
loadstore1. Thus an instruction which should generate an interrupt
such as a floating-point unavailable interrupt will instead get
executed.
To fix this, when we detect that an interrupt should be generated but
loadstore1 is still executing a previous instruction, we don't execute
any new instructions, and set a new r.intr_pending flag. That results
in busy_out being asserted (meaning that no further instructions will
come in from decode2). When loadstore1 has finished the instructions
it has, the interrupt gets sent to writeback. If one of the
instructions in loadstore1 generates an interrupt in the meantime, the
l_in.interrupt signal gets asserted and that clears r.intr_pending, so
the interrupt we detected gets discarded.
Signed-off-by: Paul Mackerras <paulus@ozlabs.org>
This implements a 1-entry partition table, so that instead of getting
the process table base address from the PRTBL SPR, the MMU now reads
the doubleword pointed to by the PTCR register plus 8 to get the
process table base address. The partition table entry is cached.
Having the PTCR and the vestigial partition table reduces the amount
of software change required in Linux for Microwatt support.
Signed-off-by: Paul Mackerras <paulus@ozlabs.org>
If the DAR and DSISR are read before they are written, we assert with:
register_file.vhdl:55:25:@60195ns:(report note): Writing GPR 09 00000000XXXXXXXX
register_file.vhdl:61:17:@60195ns:(assertion failure): Assertion violation
This initialises DAR/DSISR to avoid this.
Signed-off-by: Michael Neuling <mikey@neuling.org>
The idea here is that we can have multiple instructions in progress at
the same time as long as they all go to the same unit, because that
unit will keep them in order. If we get an instruction for a
different unit, we wait for all the previous instructions to finish
before executing it. Since the loadstore unit is the only one that is
currently pipelined, this boils down to saying that loadstore
instructions can go ahead while l_in.in_progress = 1 but other
instructions have to wait until it is 0.
This gives a 2% increase on coremark performance on the Arty A7-100
(from ~190 to ~194).
Signed-off-by: Paul Mackerras <paulus@ozlabs.org>
This makes loadstore use a 3-stage pipeline. For now, only one
instruction goes through the pipe at a time. Completion and writeback
are still combinatorial off the valid signal back from the dcache, so
performance should be the same as before. In future it should be able
to sustain one load or store per cycle provided they hit in the
dcache.
Signed-off-by: Paul Mackerras <paulus@ozlabs.org>
This fixes two bugs which show up when multiple operations are in
flight in the dcache, and adds a 'hold' input which will be needed
when loadstore1 is pipelined.
The first bug is that dcache needs to sample the data for a store on
the cycle after the store request comes in even if the store request
is held up because of a previous request (e.g. if the previous request
is a load miss or a dcbz).
The second bug is that a load request coming in for a cache line being
refilled needs to be handled immediately in the case where it is for
the row whose data arrives on the same cycle. If it is not, then it
will be handled as a separate cache miss and the cache line will be
refilled again into a different way, leading to two ways both being
valid for the same tag. This can lead to data corruption, in the
scenario where subsequent writes go to one of the ways and then that
way gets displaced but the other way doesn't. This bug could in
principle show up even without having multiple operations in flight in
the dcache.
Signed-off-by: Paul Mackerras <paulus@ozlabs.org>
This changes the way GPR hazards are detected and tracked. Instead of
having a model of the pipeline in gpr_hazard.vhdl, which has to mirror
the behaviour of the real pipeline exactly, we now assign a 2-bit tag
to each instruction and record which GSPR the instruction writes.
Subsequent instructions that need to use the GSPR get the tag number
and stall until the value with that tag is being written back to the
register file.
For now, the forwarding paths are disabled. That gives about a 8%
reduction in coremark performance.
Signed-off-by: Paul Mackerras <paulus@ozlabs.org>
This uses the instruction-doubling machinery to send load with update
instructions down to loadstore1 as two separate ops, rather than
one op with two destinations. This will help to simplify the value
tracking mechanisms.
Signed-off-by: Paul Mackerras <paulus@ozlabs.org>
This makes timing easier and also means that store floating-point
single precision instructions no longer need to take an extra cycle.
Signed-off-by: Paul Mackerras <paulus@ozlabs.org>
This makes it simpler to work out when to deliver a FPU unavailable
interrupt. This also means we can get rid of the OP_FPLOAD and
OP_FPSTORE insn_type values.
Signed-off-by: Paul Mackerras <paulus@ozlabs.org>
This implements the lq, stq, lqarx and stqcx. instructions.
These instructions all access two consecutive GPRs; for example the
"lq %r6,0(%r3)" instruction will load the doubleword at the address
in R3 into R7 and the doubleword at address R3 + 8 into R6. To cope
with having two GPR sources or destinations, the instruction gets
repeated at the decode2 stage, that is, for each lq/stq/lqarx/stqcx.
coming in from decode1, two instructions get sent out to execute1.
For these instructions, the RS or RT register gets modified on one
of the iterations by setting the LSB of the register number. In LE
mode, the first iteration uses RS|1 or RT|1 and the second iteration
uses RS or RT. In BE mode, this is done the other way around. In
order for decode2 to know what endianness is currently in use, we
pass the big_endian flag down from icache through decode1 to decode2.
This is always in sync with what execute1 is using because only rfid
or an interrupt can change MSR[LE], and those operations all cause
a flush and redirect.
There is now an extra column in the decode tables in decode1 to
indicate whether the instruction needs to be repeated. Decode1 also
enforces the rule that lq with RT = RT and lqarx with RA = RT or
RB = RT are illegal.
Decode2 now passes a 'repeat' flag and a 'second' flag to execute1,
and execute1 passes them on to loadstore1. The 'repeat' flag is set
for both iterations of a repeated instruction, and 'second' is set
on the second iteration. Execute1 does not take asynchronous or
trace interrupts on the second iteration of a repeated instruction.
Loadstore1 uses 'next_addr' for the second iteration of a repeated
load/store so that we access the second doubleword of the memory
operand. Thus loadstore1 accesses the doublewords in increasing
memory order. For 16-byte loads this means that the first iteration
writes GPR RT|1. It is possible that RA = RT|1 (this is a legal
but non-preferred form), meaning that if the memory operand was
misaligned, the first iteration would overwrite RA but then the
second iteration might take a page fault, leading to corrupted state.
To avoid that possibility, 16-byte loads in LE mode take an
alignment interrupt if the operand is not 16-byte aligned. (This
is the case anyway for lqarx, and we enforce it for lq as well.)
Signed-off-by: Paul Mackerras <paulus@ozlabs.org>
This adds code to loadstore1 to convert between single-precision and
double-precision formats, and implements the lfs* and stfs*
instructions. The conversion processes are described in Power ISA
v3.1 Book 1 sections 4.6.2 and 4.6.3.
These conversions take one cycle, so lfs* and stfs* are one cycle
slower than lfd* and stfd*.
Signed-off-by: Paul Mackerras <paulus@ozlabs.org>
This extends the register file so it can hold FPR values, and
implements the FP loads and stores that do not require conversion
between single and double precision.
We now have the FP, FE0 and FE1 bits in MSR. FP loads and stores
cause a FP unavailable interrupt if MSR[FP] = 0.
The FPU facilities are optional and their presence is controlled by
the HAS_FPU generic passed down from the top-level board file. It
defaults to true for all except the A7-35 boards.
Signed-off-by: Paul Mackerras <paulus@ozlabs.org>
Load-and-reserve and store-conditional instructions are required to
generate an alignment interrupt (0x600 vector) if their EA is not
aligned. Implement this.
Signed-off-by: Paul Mackerras <paulus@ozlabs.org>
In 32-bit mode, effective addresses are truncated to 32 bits, both for
instruction fetches and data accesses, and CR0 is set for Rc=1 (record
form) instructions based on the lower 32 bits of the result rather
than all 64 bits.
Signed-off-by: Paul Mackerras <paulus@ozlabs.org>
Big-endian mode affects both instruction fetches and data accesses.
For instruction fetches, we byte-swap each word read from memory when
writing it into the icache data RAM, and use a tag bit to indicate
whether each cache line contains instructions in BE or LE form.
For data accesses, we simply need to invert the existing byte_reverse
signal in BE mode. The only thing to be careful of is to get the sign
bit from the correct place when doing a sign-extending load that
crosses two doublewords of memory.
For now, interrupts unconditionally set MSR[LE]. We will need some
sort of interrupt-little-endian bit somewhere, perhaps in LPCR.
This also fixes a debug report statement in fetch1.vhdl.
Signed-off-by: Paul Mackerras <paulus@ozlabs.org>
This rearranges the code used for store data formatting so that the
"for i in 0 to 7" loop indexes the output bytes rather than the
input bytes. The new expression is formally identical to the old
but is easier to synthesize. This reduces the number of LUTs by
about 250 on the Artix-7 and improves timing.
Signed-off-by: Paul Mackerras <paulus@ozlabs.org>
This reworks the way that the busy and done signals are generated in
loadstore in order to work around some problems where yosys/nextpnr
are reporting combinatorial loops (not in fact on the current code but
on minor variations needed for supporting the FPU). It seems that
yosys has problems with the case statement on v.state.
This also lifts the maddr and byte_sel generation out of the case
statement. The overall result is a slight reduction in resource usage
(~30 6-input LUTs on the A7-100).
Signed-off-by: Paul Mackerras <paulus@ozlabs.org>
This computes the address sent to the MMU separately from that sent
to the dcache. This means that the address sent to the MMU doesn't
have the delay through the lsu_sum adder, making it available earlier.
The path through the lsu_sum adder and through the MMU to the MMU
done and err outputs showed up as a critical path on some builds.
Signed-off-by: Paul Mackerras <paulus@ozlabs.org>
This makes the calculation of busy as simple as possible and dependent
only on register outputs. The timing of busy is critical, as it gates
the valid signal for the next instruction, and therefore any delays
in dropping busy at the end of a load or store directly impact the
timing of a host of other paths.
This also separates the 'done without error' and 'done with error'
cases from the MMU into separate signals that are both driven directly
from registers.
Signed-off-by: Paul Mackerras <paulus@ozlabs.org>
This adds "if LOG_LENGTH > 0 generate" to the places in the core
where log output data is latched, so that when LOG_LENGTH = 0 we
don't create the logic to collect the data which won't be stored.
Signed-off-by: Paul Mackerras <paulus@ozlabs.org>
The computation of two_dwords from r.second_bytes has shown up as
part of a critical path at times. Instead we add a 'last_dword'
flag to the reg_stage_t record which tells us more directly
whether a valid flag coming in from dcache means that the
instruction is done, thereby shortening the path to the busy output
back to execute1.
This also simplifies some of the trim_ctl logic. The two_dwords = 0
case could never have use_second(i) = 1 for any of the bytes being
transferred, so "not use_second(i)" is always 1.
Signed-off-by: Paul Mackerras <paulus@ozlabs.org>
This reduces the number of cycles where loadstore1 asserts its busy
output, leading to increased throughput of loads and stores. Loads
that hit in the cache can now be executed at the rate of one every two
cycles. Stores take 4 cycles assuming the wishbone slave responds
with an ack the cycle after we assert strobe.
To achieve this, the state machine code is split into two parts, one
for when we have an existing instruction in progress, and one for
starting a new instruction. We can now combinatorially clear busy and
start a new instruction in the same cycle that we get a done signal
from the dcache; in other words we are completing one instruction and
potentially writing back results in the same cycle that we start a new
instruction and send its address and data to the dcache.
Signed-off-by: Paul Mackerras <paulus@ozlabs.org>
This makes mfspr and mtspr complete (and mfspr write back) on the
cycle after the instruction is received from execute1, rather than
on the same cycle. This makes them match all other instructions
that execute in one cycle. Because these instructions are marked
as single-issue, there wasn't the possibility of having two
instructions complete on the same cycle (which we can't cope with),
but it is better to fix this.
Signed-off-by: Paul Mackerras <paulus@ozlabs.org>
This changes the instruction dependency tracking so that we can
generate a "busy" signal from execute1 and loadstore1 which comes
along one cycle later than the current "stall" signal. This will
enable us to signal busy cycles only when we need to from loadstore1.
The "busy" signal from execute1/loadstore1 indicates "I didn't take
the thing you gave me on this cycle", as distinct from the previous
stall signal which meant "I took that but don't give me anything
next cycle". That means that decode2 proactively gives execute1
a new instruction as soon as it has taken the previous one (assuming
there is a valid instruction available from decode1), and that then
sits in decode2's output until execute1 can take it. So instructions
are issued by decode2 somewhat earlier than they used to be.
Decode2 now only signals a stall upstream when its output buffer is
full, meaning that we can fill up bubbles in the upstream pipe while a
long instruction is executing. This gives a small boost in
performance.
This also adds dependency tracking for rA updates by update-form
load/store instructions.
The GPR and CR hazard detection machinery now has one extra stage,
which may not be strictly necessary. Some of the code now really
only applies to PIPELINE_DEPTH=1.
Signed-off-by: Paul Mackerras <paulus@ozlabs.org>
This logs 256 bits of data per cycle to a ring buffer in BRAM. The
data collected can be read out through 2 new SPRs or through the
debug interface.
The new SPRs are LOG_ADDR (724) and LOG_DATA (725). LOG_ADDR contains
the buffer write pointer in the upper 32 bits (in units of entries,
i.e. 32 bytes) and the read pointer in the lower 32 bits (in units of
doublewords, i.e. 8 bytes). Reading LOG_DATA gives the doubleword
from the buffer at the read pointer and increments the read pointer.
Setting bit 31 of LOG_ADDR inhibits the trace log system from writing
to the log buffer, so the contents are stable and can be read.
There are two new debug addresses which function similarly to the
LOG_ADDR and LOG_DATA SPRs. The log is frozen while either or both of
the LOG_ADDR SPR bit 31 or the debug LOG_ADDR register bit 31 are set.
The buffer defaults to 2048 entries, i.e. 64kB. The size is set by
the LOG_LENGTH generic on the core_debug module. Software can
determine the length of the buffer because the length is ORed into the
buffer write pointer in the upper 32 bits of LOG_ADDR. Hence the
length of the buffer can be calculated as 1 << (31 - clz(LOG_ADDR)).
There is a program to format the log entries in a somewhat readable
fashion in scripts/fmt_log/fmt_log.c. The log_entry struct in that
file describes the layout of the bits in the log entries.
Signed-off-by: Paul Mackerras <paulus@ozlabs.org>
This adds the PID register and repurposes SPR 720 as the PRTBL
register, which points to the base of the process table. There
doesn't seem to be any point to implementing the partition table given
that we don't have hypervisor mode.
The MMU caches entry 0 of the process table internally (in pgtbl3)
plus the entry indexed by the value in the PID register (pgtbl0).
Both caches are invalidated by a tlbie[l] with RIC=2 or by a move to
PRTBL. The pgtbl0 cache is invalidated by a move to PID. The dTLB
and iTLB are cleared by a move to either PRTBL or PID.
Which of the two page table root pointers is used (pgtbl0 or pgtbl3)
depends on the MSB of the address being translated. Since the segment
checking ensures that address(63) = address(62), this is sufficient to
map quadrants 0 and 3.
Signed-off-by: Paul Mackerras <paulus@ozlabs.org>
Slbia (with IH=7) is used in the Linux kernel to flush the ERATs
(our iTLB/dTLB), so make it do that.
This moves the logic to work out whether to flush a single entry
or the whole TLB from dcache and icache into mmu. We now invalidate
all dTLB and iTLB entries when the AP (actual pagesize) field of
RB is non-zero on a tlbie[l], as well as when IS is non-zero.
Signed-off-by: Paul Mackerras <paulus@ozlabs.org>
This hooks up the connections so that an OP_FETCH_FAILED coming down
to loadstore1 will get sent to the MMU for it to do a radix tree walk
for the instruction address. The MMU then sends the resulting PTE to
the icache module to be installed in the iTLB. If no valid PTE can
be found, the MMU sends an error signal back to loadstore1 which sends
it on to execute1 to generate an ISI.
Signed-off-by: Paul Mackerras <paulus@ozlabs.org>
This adds a direct-mapped TLB to the icache, with 64 entries by default.
Execute1 now sends a "virt_mode" signal from MSR[IR] to fetch1 along
with redirects to indicate whether instruction addresses should be
translated through the TLB, and fetch1 sends that on to icache.
Similarly a "priv_mode" signal is sent to indicate the privilege
mode for instruction fetches. This means that changes to MSR[IR]
or MSR[PR] don't take effect until the next redirect, meaning an
isync, rfid, branch, etc.
The icache uses a hash of the effective address (i.e. next instruction
address) to index the TLB. The hash is an XOR of three fields of the
address; with a 64-entry TLB, the fields are bits 12--17, 18--23 and
24--29 of the address. TLB invalidations simply invalidate the
indexed TLB entry without checking the contents.
If the icache detects a TLB miss with virt_mode=1, it will send a
fetch_failed indication through fetch2 to decode1, which will turn it
into a special OP_FETCH_FAILED opcode with unit=LDST. That will get
sent down to loadstore1 which will currently just raise a Instruction
Storage Interrupt (0x400) exception.
One bit in the PTE obtained from the TLB is used to check whether an
instruction access is allowed -- the privilege bit (bit 3). If bit 3
is 1 and priv_mode=0, then a fetch_failed indication is sent down to
fetch2 and to decode1, which generates an OP_FETCH_FAILED. Any PTEs
with PTE bit 0 (EAA[3]) clear or bit 8 (R) clear should not be put
into the iTLB since such PTEs would not allow execution by any
context.
Tlbie operations get sent from mmu to icache over a new connection.
Unfortunately the privileged instruction tests are broken for now.
Signed-off-by: Paul Mackerras <paulus@ozlabs.org>
This is required by the architecture. It means that the error bits
reported in DSISR or SRR1 now come from the permission/RC check done
on the refetched PTE rather than the TLB entry. Unfortunately that
somewhat breaks the software-loaded TLB mode of operation in that
DSISR/SRR1 always report no PTE rather than permission error or
RC failure.
This also restructures the loadstore1 state machine a bit, combining
the FIRST_ACK_WAIT and LAST_ACK_WAIT states into a single state and
the MMU_LOOKUP_1ST and MMU_LOOKUP_LAST states likewise. We now have a
'dwords_done' bit to say whether the first transfer of two (for an
unaligned access) has been done.
The cache paradox error (where a non-cacheable access finds a hit in
the cache) is now the only cause of DSI from the dcache. This should
probably be a machine check rather than DSI in fact.
Signed-off-by: Paul Mackerras <paulus@ozlabs.org>
A data segment interrupt (DSegI) occurs when an address to be
translated by the MMU is outside the range of the radix tree
or the top two bits of the address (the quadrant) are 01 or 10.
This is detected in a new state of the MMU state machine, and
is sent back to loadstore1 as an error, which sends it on to
execute1 to generate an interrupt to the 0x380 vector.
Signed-off-by: Paul Mackerras <paulus@ozlabs.org>
This adds the necessary machinery to the MMU for it to do radix page
table walks. The core elements are a shifter that can shift the
address right by between 0 and 47 bits, a mask generator that can
generate a mask of between 5 and 16 bits, a final mask generator,
and new states in the state machine.
(The final mask generator is used for transferring bits of the
original address into the resulting TLB entry when the leaf PTE
corresponds to a page size larger than 4kB.)
The hardware does not implement a partition table or a process table.
Software is expected to load the appropriate process table entry
into a new SPR called PGTBL0, SPR 720. The contents should be
formatted as described in Book III section 5.7.6.2 of the Power ISA
v3.0B. PGTBL0 is set to 0 on hard reset. At present, the top two bits
of the address (the quadrant) are ignored.
There is currently no caching of any step in the translation process
or of the final result, other than the entry created in the dTLB.
That entry is a 4k page entry even if the leaf PTE found in the walk
corresponds to a larger page size.
This implementation can handle almost any page table layout and any
page size. The RTS field (in PGTBL0) can have any value between 0
and 31, corresponding to a total address space size between 2^31
and 2^62 bytes. The RPDS field of PGTBL0 can be any value between
5 and 16, except that a value of 0 is taken to disable radix page
table walking (for use when one is using software loading of TLB
entries). The NLS field of the page directory entries can have any
value between 5 and 16. The minimum page size is 4kB, meaning that
the sum of RPDS and the NLS values of the PDEs found on the path to
a leaf PTE must be less than or equal to RTS + 31 - 12.
The PGTBL0 SPR is in the mmu module; thus this adds a path for
loadstore1 to read and write SPRs in mmu. This adds code in dcache
to service doubleword read requests from the MMU, as well as requests
to write dTLB entries.
Signed-off-by: Paul Mackerras <paulus@ozlabs.org>
This adds a new module to implement an MMU. At the moment it doesn't
do very much. Tlbie instructions now get sent by loadstore1 to mmu,
which sends them to dcache, rather than loadstore1 sending them
directly to dcache. TLB misses from dcache now get sent by loadstore1
to mmu, which currently just returns an error. Loadstore1 then
generates a DSI in response to the error return from mmu.
Signed-off-by: Paul Mackerras <paulus@ozlabs.org>
This adds logic to the dcache to check the permissions encoded in
the PTE that it gets from the dTLB. The bits that are checked are:
R must be 1
C must be 1 for a store
EAA(0) - if this is 1, MSR[PR] must be 0
EAA(2) must be 1 for a store
EAA(1) | EAA(2) must be 1 for a load
In addition, ATT(0) is used to indicate a cache-inhibited access.
This now implements DSISR bits 36, 38 and 45.
(Bit numbers above correspond to the ISA, i.e. using big-endian
numbering.)
MSR[PR] is now conveyed to loadstore1 for use in permission checking.
Signed-off-by: Paul Mackerras <paulus@ozlabs.org>
This adds a path from loadstore1 back to execute1 for reporting
errors, and machinery in execute1 for generating data storage
interrupts at vector 0x300.
If dcache is given two requests in successive cycles and the
first encounters an error (e.g. a TLB miss), it will now cancel
the second request.
Loadstore1 now responds to errors reported by dcache by sending
an exception signal to execute1 and returning to the idle state.
Execute1 then writes SRR0 and SRR1 and jumps to the 0x300 Data
Storage Interrupt vector. DAR and DSISR are held in loadstore1.
Signed-off-by: Paul Mackerras <paulus@ozlabs.org>
This adds a TLB to dcache, providing the ability to translate
addresses for loads and stores. No protection mechanism has been
implemented yet. The MSR_DR bit controls whether addresses are
translated through the TLB.
The TLB is a fixed-pagesize, set-associative cache. Currently
the page size is 4kB and the TLB is 2-way set associative with 64
entries per set.
This implements the tlbie instruction. RB bits 10 and 11 control
whether the whole TLB is invalidated (if either bit is 1) or just
a single entry corresponding to the effective page number in bits
12-63 of RB.
As an extension until we get a hardware page table walk, a tlbie
instruction with RB bits 9-11 set to 001 will load an entry into
the TLB. The TLB entry value is in RS in the format of a radix PTE.
Currently there is no proper handling of TLB misses. The load or
store will not be performed but no interrupt is generated.
In order to make timing at 100MHz on the Arty A7-100, we compare
the real address from each way of the TLB with the tag from each way
of the cache in parallel (requiring # TLB ways * # cache ways
comparators). Then the result is selected based on which way hit in
the TLB. That avoids a timing path going through the TLB EA
comparators, the multiplexer that selects the RA, and the cache tag
comparators.
The hack where addresses of the form 0xc------- are marked as
cache-inhibited is kept for now but restricted to real-mode accesses.
Signed-off-by: Paul Mackerras <paulus@ozlabs.org>
This arranges for some mfspr and mtspr to get sent to loadstore1
instead of being handled in execute1. In particular, DAR and DSISR
are handled this way. They are therefore "slow" SPRs.
While we're at it, fix the spelling of HEIR and remove mention of
DAR and DSISR from the comments in execute1.
Signed-off-by: Paul Mackerras <paulus@ozlabs.org>
This adds logic to dcache and loadstore1 to implement dcbz. For now
it zeroes a single cache line (by default 64 bytes), not 128 bytes
like IBM Power processors do.
The dcbz operation is performed much like a load miss, except that
we are writing zeroes to memory instead of reading. As each ack
comes back, we write zeroes to the BRAM instead of data from memory.
In this way we zero the line in memory and also zero the line of
cache memory, establishing the line in the cache if it wasn't already
resident. If it was already resident then we overwrite the existing
line in the cache.
Signed-off-by: Paul Mackerras <paulus@ozlabs.org>
In preparation for adding a TLB to the dcache, this plumbs the
insn_type from execute1 through to loadstore1, so that we can have
other operations besides loads and stores (e.g. tlbie) going to
loadstore1 and thence to the dcache. This also plumbs the unit field
of the decode ROM from decode2 through to execute1 to simplify the
logic around which ops need to go to loadstore1.
The load and store data formatting are now not conditional on the
op being OP_LOAD or OP_STORE. This eliminates the inferred latches
clocked by each of the bits of r.op that we were getting previously.
Signed-off-by: Paul Mackerras <paulus@ozlabs.org>
This adds support for lbzcix, lhzcix, lwzcix, ldcix, stbcix, sthcix,
stwcix and stdcix. The temporary hack where accesses to addresses of
the form 0xc??????? are made non-cacheable is left in for now to avoid
making existing programs non-functional.
Signed-off-by: Paul Mackerras <paulus@ozlabs.org>
This puts all the data formatting (byte rotation based on lowest three
bits of the address, byte reversal, sign extension, zero extension)
in loadstore1. Writeback now simply sends the data provided to the
register files.
Signed-off-by: Paul Mackerras <paulus@ozlabs.org>
So that the dcache could in future be used by an MMU, this moves
logic to do with data formatting, rA updates for update-form
instructions, and handling of unaligned loads and stores out of
dcache and into loadstore1. For now, dcache connects only to
loadstore1, and loadstore1 now has the connection to writeback.
Dcache generates a stall signal to loadstore1 which indicates that
the request presented in the current cycle was not accepted and
should be presented again. However, loadstore1 doesn't currently
use it because we know that we can never hit the circumstances
where it might be set.
For unaligned transfers, loadstore1 generates two requests to
dcache back-to-back, and then waits to see two acks back from
dcache (cycles where d_in.valid is true).
Loadstore1 now has a FSM for tracking how many acks we are
expecting from dcache and for doing the rA update cycles when
necessary. Handling for reservations and conditional stores is
still in dcache.
Loadstore1 now generates its own stall signal back to decode2,
so we no longer need the logic in execute1 that generated the stall
for the first two cycles.
Signed-off-by: Paul Mackerras <paulus@ozlabs.org>
Currently we don't get the result from a load that hits in the dcache
until the fourth cycle after the instruction was presented to
loadstore1. This trims this back to 3 cycles by taking the low order
bits of the address generated in loadstore1 into dcache directly (not
via the output register of loadstore1) and using them to address the
read port of the dcache data RAM. We use the lower 12 address bits
here in the expectation that any reasonable data cache design will
have a set size of 4kB or less in order to avoid the aliasing problems
that can arise with a virtually-indexed physically-tagged cache if
the set size is greater than the smallest page size provided by the
MMU.
With this we can get rid of r2 and drive the signals going to
writeback from r1, since the load hit data is now available one
cycle earlier. We need a multiplexer on the read address of the
data cache RAM in order to handle the second doubleword of an
unaligned access.
One small complication is that we now need an extra cycle in the case
of an unaligned load which misses in the data cache and which reads
the 2nd-last and last doublewords of a cache line. This is the reason
for the PRE_NEXT_DWORD state; if we just go straight to NEXT_DWORD
then we end up having the write of the last doubleword of the cache
line and the read of that same doubleword occurring in the same
cycle, which means we read stale data rather than the just-fetched
data.
Signed-off-by: Paul Mackerras <paulus@ozlabs.org>
