This adds an 'is_signed' signal to MultiplyInputType to indicate
whether the data1 and data2 fields are to be interpreted as signed or
unsigned numbers.
The 'not_result' field is replaced by a 'subtract' field which
provides a more intuitive interface for requesting that the product be
subtracted from the addend rather than added, i.e. subtract = 1 gives
C - A * B, vs. subtract = 0 giving C + A * B. (Previously the users
of the multipliers got the same effect by complementing the addend and
setting not_result = 1.)
The is_32bit field is removed because it is no longer used now that we
have a separate 32-bit multiplier.
Signed-off-by: Paul Mackerras <paulus@ozlabs.org>
This splits out the decoding done in the decode0 step into a separate
predecoder, used when writing instructions into the icache. The
icache now holds 36 bits per instruction rather than 32. For valid
instructions, those 36 bits comprise the bottom 26 bits of the
instruction word, a 9-bit insn_code value (which uniquely identifies
the instruction), and a zero in the MSB. For illegal instructions,
the MSB is one and the full instruction word is in the bottom 32 bits.
Having the full instruction word available for illegal instructions
means that it can be printed in the log when simulating, or in future
could be placed in the HEIR register.
If we don't have an FPU, then the floating-point instructions are
regarded as illegal. In that case, the insn_code values would fit
into 8 bits, which could be used in future to reduce the size of
decode_rom from 512 to 256 entries.
Signed-off-by: Paul Mackerras <paulus@ozlabs.org>
This reduces the block RAM requirements for instruction decoding by
splitting it into two steps. The first, in a new pipeline stage
called decode0 (implemented by code in decode1.vhdl) maps the
instruction to a 9-bit instruction code using major and row decode
ROMs. The second maps the 9-bit code to the final decode_rom_t (about
44 bits wide). Branch prediction done in decode is now done in
decode0 rather than decode1.
Signed-off-by: Paul Mackerras <paulus@ozlabs.org>
Do more decoding of the instruction ahead of the IDLE state
processing so that the IDLE state code becomes much simpler.
To make the decoding easier, we now use four insn_type_t codes for
floating-point operations rather than two. This also rearranges the
insn_type_t values a little to get the 4 FP opcode values to differ
only in the bottom 2 bits, and put OP_DIV, OP_DIVE and OP_MOD next to
them.
Signed-off-by: Paul Mackerras <paulus@ozlabs.org>
Instead of having a multiplexer in loadstore1 in order to be able to
put the instruction address into v.addr, we now set decode.input_reg_a
to CIA in the decode table entry for OP_FETCH_FAILED. That means that
the operand selection machinery in decode2 will supply the instruction
address to loadstore1 on the lv.addr1 input and no special case is
needed in loadstore1. This saves a few LUTs (~40 on the Artix-7).
Signed-off-by: Paul Mackerras <paulus@ozlabs.org>
This removes some logic that was previously added for the 16-byte
loads and stores (lq, lqarx, stq, stqcx.) and not completely removed
in commit c9e838b656 ("Remove support for lq, stq, lqarx and
stqcx.", 2022-06-04).
Signed-off-by: Paul Mackerras <paulus@ozlabs.org>
This improves timing a little because the register addresses now come
directly from a latch instead of being calculated by
decode_input_reg_*. The asserts that check that the two are the same
are now in decode2 rather than register_file.
Signed-off-by: Paul Mackerras <paulus@ozlabs.org>
With this, the register RAM is read synchronously using the addresses
supplied by decode1. That means the register RAM can now be block RAM
rather than LUT RAM.
Debug accesses are done via the B port on cycles when decode1
indicates that there is no valid instruction or the instruction
doesn't use a [F]RB operand.
We latch the addresses being read in each cycle and use the same
address next cycle if stalled. Data that is being written is latched
and a multiplexer on each read port then supplies the latched write
data if the read address for that port equals the write address.
Signed-off-by: Paul Mackerras <paulus@ozlabs.org>
This adds some relatively simple logic to decode1 to compute the
GPR/FPR addresses that an instruction will access. It always computes
three addresses regardless of whether the instruction will actually
use all of them. The main things it computes are whether the
instruction uses the RS field or the RC field for the 3rd operand, and
whether the operands are FPRs or GPRs (it is possible for RS to be an
FPR but RA and RB to be GPRs, as for example with stfdx).
At the moment all we do with these computed register addresses is to
assert that they are identical to the ones coming from decode2 one
cycle later.
Signed-off-by: Paul Mackerras <paulus@ozlabs.org>
This provides access to the SPRs via the JTAG DMI interface. For now
they are still accessed as if they were GPR/FPRs using the same
numbering as before (GPRs at 0 - 0x1f, SPRs at 0x20 - 0x2d, FPRs at
0x40 - 0x5f).
For XER, debug reads now report the full value, not just the bits that
were previously stored in the register file. The "slow" SPR mux is
not used for debug reads.
Decode2 determines on each cycle whether a debug SPR access will
happen next cycle, based on whether there is a request and whether the
current instruction accesses the SPR RAM.
Signed-off-by: Paul Mackerras <paulus@ozlabs.org>
With this, the register file now contains 64 entries, for 32 GPRs and
32 FPRs, rather than the 128 it had previously. Several things get
simplified - decode1 no longer has to work out the ispr{1,2,o} values,
decode_input_reg_{a,b,c} no longer have the t = SPR case, etc.
Signed-off-by: Paul Mackerras <paulus@ozlabs.org>
By putting CTR on the odd side and LR and TAR on the even side, we can
read and write CTR for bdnz-style instructions in parallel with
reading LR or TAR for indirect branches and writing LR for branches
with LK=1. Thus we don't need to double up any of these instructions,
giving a simplification in decode2.
We now have logic for printing LR and CTR at the end of a simulation
in execute1, in addition to the similar logic in register_file and
cr_file.
Signed-off-by: Paul Mackerras <paulus@ozlabs.org>
This starts the process of removing SPRs from the register file by
moving SRR0/1, SPRG0-3, HSRR0/1 and HSPRG0/1 out of the register file
and putting them into execute1. They are stored in a pair of small
RAM arrays, referred to as "even" and "odd". The reason for having
two arrays is so that two values can be read and written in each
cycle. For example, SRR0 and SRR1 can be written in parallel by an
interrupt and read in parallel by the rfid instruction.
The addresses in the RAM which will be accessed are determined in the
decode2 stage. We have one write address for both sides, but two read
addresses, since in future we will want to be able to read CTR at the
same time as either LR or TAR.
We now have a connection from writeback to execute1 which carries the
partial SRR1 value for an interrupt. SRR0 comes from the execute
pipeline; we no longer need to carry instruction addresses along the
LSU and FPU pipelines. Since SRR0 and SRR1 can be written in the same
cycle now, we don't need the little state machine in writeback any
more.
Signed-off-by: Paul Mackerras <paulus@ozlabs.org>
- Arrange for XER to be written for OE=1 forms
- Arrange for condition codes to be set for RC=1 forms
(including correct handling for 32-bit mode)
- Don't instantiate the divider if we have an FPU.
Signed-off-by: Paul Mackerras <paulus@ozlabs.org>
This adds logic to the FPU to accomplish 64-bit integer divisions.
No instruction actually uses this yet.
The algorithm used is to obtain an estimate of the reciprocal of the
divisor using the lookup table and refine it by one to three
iterations of the Newton-Raphson algorithm (the number of iterations
depends on the number of significant bits in the dividend). Then the
reciprocal is multiplied by the dividend to get the quotient estimate.
The remainder is calculated as dividend - quotient * divisor. If the
remainder is greater than or equal to the divisor, the quotient is
incremented, or if a modulo operation is being done, the divisor is
subtracted from the remainder. The inverse estimate after refinement
is good enough that the quotient estimate is always equal to or one
less than the true quotient.
Signed-off-by: Paul Mackerras <paulus@ozlabs.org>
At present the busy/stall signal going to decode1 depends on whether
control thinks it can issue the current instruction, and that depends
on completion and bypass signals coming from execute1 and writeback.
To improve the timing of stall_out, this rearranges decode2 so that
stall_out is asserted when we have a valid instruction that couldn't
be issued in the previous cycle. This means that decode1 could give
us a new instruction when we haven't issued the previous instruction.
This in turn means that we can only use d_in in the first cycle of
processing an instruction. After the first cycle, we get register
addresses etc. from dc2 rather than d_in.
Then, to avoid the need to read register operands from register_file
in each cycle until the instruction issues, we bring the bypass path
for data being written to the register file into decode2 explicitly
rather than having it in register_file.
A new process called decode2_addrs does the process of calling
decode_input_reg_* and decode_output_reg and sets up the register file
addresses. This was split out (and decode_input_reg_* reworked) to
try to reduce the number of passes through the decode2_1 process that
need to be done in simulation.
Signed-off-by: Paul Mackerras <paulus@ozlabs.org>
This makes the FPU able to stall other units at execute stage 2 and be
stalled by other units (specifically the LSU).
This means that the completion and writeback for an instruction can
now end up being deferred until the second cycle of a following
instruction, i.e. the cycle when the state machine has gone through
IDLE state into one of the DO_* states, which means we need to latch
the destination FPR number, CR mask, etc. from the previous
instruction so that we present the correct information to writeback.
The advantage of this is that we can get rid of the in_progress signal
from the LSU.
Signed-off-by: Paul Mackerras <paulus@ozlabs.org>
Execute1 and loadstore1 now send each other stall signals that
indicate that a valid instruction in stage 2 can't complete in this
cycle, and hence any valid instruction in stage 1 in the other unit
can't move to stage 2. With this in place, an ALU instruction can
move into stage 1 while a LSU instruction is in stage 2.
Since the FPU doesn't yet have a way to stall completion, we can't yet
start FPU instructions while any LSU or ALU instruction is in
progress.
Signed-off-by: Paul Mackerras <paulus@ozlabs.org>
This adds a second execute stage to the pipeline, in order to match up
the length of the pipeline through loadstore and dcache with the
length through execute1. This will ultimately enable us to get rid of
the 1-cycle bubble that we currently have when issuing ALU
instructions after one or more LSU instructions.
Most ALU instructions execute in the first stage, except for
count-zeroes and popcount instructions (which take two cycles and do
some of their work in the second stage) and mfspr/mtspr to "slow" SPRs
(TB, DEC, PVR, LOGA/LOGD, CFAR). Multiply and divide/mod instructions
take several cycles but the instruction stays in the first stage (ex1)
and ex1.busy is asserted until the operation is complete.
There is currently a bypass from the first stage but not the second
stage. Performance is down somewhat because of that and because this
doesn't yet eliminate the bubble between LSU and ALU instructions.
The forwarding of XER common bits has been changed somewhat because
now there is another pipeline stage between ex1 and the committed
state in cr_file. The simplest thing for now is to record the last
value written and use that, unless there has been a flush, in which
case the committed state (obtained via e_in.xerc) is used.
Note that this fixes what was previously a benign bug in control.vhdl,
where it was possible for control to forget an instructions dependency
on a value from a previous instruction (a GPR or the CR) if this
instruction writes the value and the instruction gets to the point
where it could issue but is blocked by the busy signal from execute1.
In that situation, control may incorrectly not indicate that a bypass
should be used. That didn't matter previously because, for ALU and
FPU instructions, there was only one previous instruction in flight
and once the current instruction could issue, the previous instruction
was completing and the correct value would be obtained from
register_file or cr_file. For loadstore instructions there could be
two being executed, but because there are no bypass paths, failing to
indicate use of a bypass path is fine.
Signed-off-by: Paul Mackerras <paulus@ozlabs.org>
We now have a record that represents the actions taken in executing an
instruction, and a process that computes that for the incoming
instruction. We no longer have 'current' or 'r.cur_instr', instead
things like the destination register are put into r.e in the first
cycle of an instruction and not reinitialized in subsequent busy
cycles.
For mfspr and mtspr, we now decode "slow" SPR numbers (those SPRs that
are not stored in the register file) to a new "spr_selector" record
in decode1 (excluding those in the loadstore unit). With this, the
result for mfspr is determined in the data path.
Signed-off-by: Paul Mackerras <paulus@ozlabs.org>
Besides the overflow and status carry bits, XER has 18 bits which need
to retain the value written by mtxer (in case software wants to
emulate the move-assist instructions (lswi, lswx, stswi, stswx).
Until now these bits (and others) have been stored in the GPR file as
a "fast" SPR, but this causes complications because XER is not really
a fast SPR.
Instead, we now store these 18 bits in the 'ctrl' signal, which exists
in execute1. This will enable us to simplify the data path in future,
and has the added bonus that with a little bit of plumbing, we can get
the full XER value printed when dumping registers at the end of a
simulation.
Therefore this changes scripts/run_test.sh to remove the greps which
exclude XER from the comparison of actual and expected register
results.
Signed-off-by: Paul Mackerras <paulus@ozlabs.org>
This moves REAL_ADDR_BITS out of the caches and defines a real_addr_t
type for a real address, along with a addr_to_real() conversion helper.
It makes the vhdl a bit more readable
Signed-off-by: Benjamin Herrenschmidt <benh@kernel.crashing.org>
This adds a bit to the BTC to store whether the corresponding branch
instruction was taken last time it was encountered. That lets us pass
a not-taken prediction down to decode1, which for backwards direct
branches inhibits it from redirecting fetch to the target of the
branch. This increases coremark by about 2%.
Signed-off-by: Paul Mackerras <paulus@ozlabs.org>
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>
This is the start of an implementation of a PMU according to PowerISA
v3.0B. Things not implemented yet include most architected events,
the BHRB, event-based branches, thresholding, MMCR0[TBCC] field, etc.
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>
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 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 moves the logic for redirecting fetching and writing SRR0 and
SRR1 to writeback. The aim is that ultimately units other than
execute1 can send their interrupts to writeback along with their
instruction completions, so that there can be multiple instructions
in flight without needing execute1 to keep track of the address
of each outstanding instruction.
Signed-off-by: Paul Mackerras <paulus@ozlabs.org>
This changes the bypass path. Previously it went from after
execute1's output to after decode2's output. Now it goes from before
execute1's output register to before decode2's output register. The
reason is that the new path will be simpler to manage when there are
possibly multiple instructions in flight. This means that the
bypassing can be managed inside decode2 and control.
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 convert conditional
branch instructions that update both CTR and LR (e.g., bdnzl, bdnzlrl)
into two instructions, of which the first updates CTR and determines
whether the branch is taken, and the second updates LR and does the
redirect if necessary.
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 implements a cache in fetch1, where each entry stores the address
of a simple branch instruction (b or bc) and the target of the branch.
When fetching sequentially, if the address being fetched matches the
cache entry, then fetching will be redirected to the branch target.
The cache has 1024 entries and is direct-mapped, i.e. indexed by bits
11..2 of the NIA.
The bus from execute1 now carries information about taken and
not-taken simple branches, which fetch1 uses to update the cache.
The cache entry is updated for both taken and not-taken branches, with
the valid bit being set if the branch was taken and cleared if the
branch was not taken.
If fetching is redirected to the branch target then that goes down the
pipe as a predicted-taken branch, and decode1 does not do any static
branch prediction. If fetching is not redirected, then the next
instruction goes down the pipe as normal and decode1 does its static
branch prediction.
In order to make timing, the lookup of the cache is pipelined, so on
each cycle the cache entry for the current NIA + 8 is read. This
means that after a redirect (from decode1 or execute1), only the third
and subsequent sequentially-fetched instructions will be able to be
predicted.
This improves the coremark value on the Arty A7-100 from about 180 to
about 190 (more than 5%).
The BTC is optional. Builds for the Artix 7 35-T part have it off by
default because the extra ~1420 LUTs it takes mean that the design
doesn't fit on the Arty A7-35 board.
Signed-off-by: Paul Mackerras <paulus@ozlabs.org>
This breaks up the enormous if .. elsif .. case .. elsif statement in
execute1 in order to try to make it simpler and more understandable.
We now have decode2 deciding whether the instruction has a value to be
written back to a register (GPR, GSPR, FPR, etc.) rather than
individual cases in execute1 setting result_en. The computation of
the data to be written back is now independent of detection of various
exception conditions. We now have an if block determining if any
exception condition exists which prevents the next instruction from
being executed, then the case statement which performs actions such as
setting carry/overflow bits, determining if a trap exception exists,
doing branches, etc., then an if statement for all the r.busy = 1
cases (continuing execution of an instruction which was started in a
previous cycle, or writing SRR1 for an interrupt).
Signed-off-by: Paul Mackerras <paulus@ozlabs.org>
This adds an explicit multiplexer feeding v.e.write_data in execute1,
with the select lines determined in the previous cycle based on the
insn_type. Similarly, for multiply and divide instructions, there is
now an explicit multiplexer.
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>
Some of the bits in the FPU buses end up as z state. Yosys
flags them, so we may as well clean it up.
Signed-off-by: Anton Blanchard <anton@linux.ibm.com>
This adds the skeleton of a floating-point unit and implements the
mffs and mtfsf instructions.
Execute1 sends FP instructions to the FPU and receives busy,
exception, FP interrupt and illegal interrupt signals from it.
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>
Anton Blanchard