CALL saves A, B and Data Pointers 0 to 2 and nothing else, which is exactly
why Q and DP3 are how a subroutine hands something back. An interrupt saves
all of it, so a service with an answer had to reach into its own frame and
un-save two fields by hand:
MVSD.2
DPUP.2 0d02 ; the saved Q, by an offset it had to know
STA.2
RETI
Thirty places in CosmOS did that. Every one knew the frame's layout by
heart, and all thirty would have gone quietly wrong the day the frame
gained a field - the same duplicated fact this project keeps being bitten
by, except duplicated into thirty places AND into the CPU.
SRET is 0x76, in the seat the block split left for it. It is RETI's frame
with RET's rule applied: A, B and DP0 to DP2 come back, the saved Q and DP3
are dropped, and the Interrupt Flag is restored from the frame - only that
bit, so carry survives a service the way it survives a call, and there is
one rule rather than two. RETI stays exactly as it was: a hardware handler
has nothing to say and must leave no trace.
CosmOS is 10,969 bytes against 11,122, and no handler knows a frame offset.
TWO MISTAKES WORTH RECORDING, both mine, both caught by tests.
The first conversion matched STA.2 with a regular expression that did not
allow a trailing comment, so it ran past the end of one handler and into
the next. The second understood the pattern and still got it wrong: the old
frame write carried the answer from A into the saved Q slot, so simply
deleting the write left Q holding whatever it happened to hold. Services
that answer by calling something were fine - Q already had it - and
services that set A directly silently reported success for every failure.
cosmosCwd is what noticed, by saying "cannot go there" about a directory
that was there. Sixteen handlers move the answer into Q now.
Seven MVQA went with it. They copied Q into A so the frame write could
carry it; SRET puts A back, so they moved a value nobody would ever read.
Three blocks move and nothing else changes. Branches take 0x60, subroutines
take 0x70, and the ALU moves up into the 0x10 block the two of them used to
share. Order within each block is preserved exactly - this relocates them,
it does not rethink them.
WHAT IT BUYS IS AN EMPTY 0x00 TO 0x0F. Program Memory that was never
written, or a load that stopped part way and left zeroes in its tail, used
to read as a long run of ADDs: the machine carried on through them, arrived
somewhere unpredictable, and whatever broke there was a long way from the
byte that caused it. Now it faults where it is met:
Fault: 0x00 at Program Address 0x0004 is not an instruction.
That is the address of the byte after the last real instruction, which is
the difference between a diagnosis and a search. Reserving the whole nibble
rather than just 0x00 means a run into blank memory faults wherever it
starts rather than only when it lands on the right byte. runOffTest records
it, and the block is left empty for whatever turns out to want it.
The other half is room: branches and subroutines had filled 0x10 to 0x1F
between them, so a service return that keeps Q and DP3 had nowhere to sit
next to its family. It has 0x76 waiting now.
Five places wrote an opcode down that the scripted remap did not reach, and
four of them were found by tests rather than by looking:
- secondPass.c lists which opcodes take an address, and firstPass.c knows
SWI by number. Missing those made XOR read as a branch.
- Asm.asm knows SWI by number too, being the other assembler. Missing it
made the native and host assemblers disagree byte for byte, which is
exactly the check that exists to catch a thing known in two places.
- loaderTest.asm carries a hand written payload, and its RETI was 0x19. To
the assembler those are numbers and to the program they are data, so
nothing but running it could notice. It says so in a comment now.
- The Assembler Manual prints the bytes hello.asm assembles to, and two of
them were branches.
The monitor's recorded disassembly moved by exactly the bytes it should:
18 became 72 wherever SWI appears, with SETD and INIB untouched and every
disassembled line still reading the same.
The loader marks the superblock before it hands over and the system clears
the mark when it reaches its prompt, so a system that crashes on the way
there leaves it set. The loader finding it still set next time is how a
machine that will not start says so to the only thing in a position to do
anything about it. Without that, pointing boot.cfg at something that dies
before the shell is a machine that can never be told anything again - the
shell is the only way to change the file, and the file is what stops the
shell from starting.
Three states rather than two, and the third is the one worth having:
0 settled the last start arrived; use the configuration
1 trying handed over, and nothing came back to say it got there
2 fell back a try failed and the fallback was used, until settled
With only 0 and 1 the machine alternates for ever: fall back, reach a
prompt, clear the mark, retry the broken system, crash, fall back. State 2
stops that. A system known not to start is not tried again until somebody
says the situation has changed.
REACHING THE PROMPT IS A DELIBERATE THRESHOLD. It is not a claim that the
system works - a shell can be reached by something broken in every other
way. It is the point where a person can type, which is exactly what the
fallback exists to give back: anything wrong past there is fixable from the
prompt and nothing wrong before it is fixable at all.
The routines live in sbfs.asm because both the loader and the system read
and write this byte, and two pieces of code with their own idea of where a
byte lives is what this format has two implementations and a byte for byte
comparison to avoid.
And the trap this system documents in its own manual caught me anyway: the
first version handed the state back in A, which CALL restores, so every
read got whatever the caller happened to be holding. It comes back in
memory now, and the comment says why.
Three disks differing only in the state on them, so the tests read as three
consecutive starts of one machine while none depends on another running.
./SplitBit --disk system.img
stage two
CosmOS
>
No boot image named. The emulator shadows its built in stage one into
Program Memory - boot vector included - and the CPU then does exactly what
it has always done: reads the boot vector and starts where it points.
NOTHING ABOUT THE CPU CHANGED to make a machine that starts itself, which
is what picking shadowing over a mapped ROM bought.
The ROM is generated from Programs/Boot/stage1.asm by the makefile rather
than committed beside it, because a copy of a program kept next to the
program is a copy that goes stale. That makes the assembler a real
dependency of the emulator, which it always sort of was and now says so.
od and awk rather than xxd, which is not everywhere, or python, which the
README does not ask anybody to install in order to build this.
loadROM is loadFile given bytes instead of a path: both go through one
reader over an fmemopen stream, because a ROM is a boot image and there is
no reason for the machine to have two ways of understanding one.
Naming an image still works and is what every other test here does. That
path is not a shortcut to apologise for - placing memory from outside is a
real thing real machines allow, and it is a debugger. The help says so now.
No image and no disk is the one case with nothing to run, and it says that
rather than printing a usage message about a missing file.
run.sh gained a "rom" mode which hands the emulator a disk and nothing
else. The source column still names stage1.asm, because that is what is IN
the ROM: assembling it there says the thing the emulator carries is a thing
that still assembles.
Stage one exists and works. It is 330 bytes of program and everything it
knows is a thing that will be true forever: which port the disk is on, that
a SplitBit disk begins with its own name, and where two numbers sit in that
first block. Not what a file is, not what a directory is, not that SBFS has
versions. It reads the live boot slot into Program Memory, jumps to the
first byte, and prints one character and halts if there is nothing there.
It is an ordinary boot image for now, so the whole chain runs on machinery
that already exists and the emulator has not been touched. Nothing about it
changes when it moves into ROM except who puts it in memory.
SplitDisk gained "boot" to write a slot and "bootslot" to choose one, kept
apart on purpose: writing a slot and starting from it are different
decisions, and joining them would make every write a commitment. A slot is
always written WHOLE, because one still holding the tail of what was there
before is one whose contents depend on its history, and stage one reads all
of it without knowing where the file stopped.
Three recorded tests, and the pair is the point: two disks differing only
in which slot the superblock names, with payloads that say different
things. One prints "booted" and the other does not, so this is a test of
CHOOSING a slot rather than a test that some bytes were read. The third
boots a disk with no boot area and gets the one character a ROM has room
for. Eight more host checks, including that a slot is padded whole.
Two things worth recording. The first draft used #Align to put the scratch
buffer at 0x8000 and produced a 33K file - thirty two kilobytes of zeroes
in something meant to be a ROM. It is an address, not storage, which is
exactly what the assembler's own scratch map exists to say.
And SplitLint caught the second in code written an hour after the baseline
that catches it. In the blit set-up, RSTA writes a source address of zero
and then RSTA writes a bank number of zero - two unrelated quantities that
are equal by accident, in the most safety critical file in the repository.
It is marked with a reason rather than removed.
The first rung of booting from disk. A boot area is blocks between the
superblock and the directory that the filesystem never allocates and never
sees, and NOTHING WAS ADDED TO RESERVE THEM: both implementations work out
the first usable block as directoryStart + directoryBlocks, and
directoryStart has always been a field rather than a constant. Formatting
with the directory moved up reserves everything below it. Neither allocator
changed, on either side.
Two new superblock fields in bytes that were reserved: bootBlocks at 14,
per slot, and bootSlot at 16. A disk made before this has zero in both,
which reads as "no boot area" - true, and the same shape as the version two
parent field, where the value an older disk already held was the right
answer without conversion.
TWO SLOTS, ALWAYS. A boot slot is raw blocks with no entry to rename, so
the write-a-temporary-and-rename ordering that protects every file cannot
protect it, and a machine interrupted while updating its only slot would
not boot at all - the one failure on this disk with no way back. Writing
the slot that is not live and then moving one byte makes that a machine
that boots what it had before.
bootBlocks and directoryStart say the same thing from two sides, so a disk
where they disagree is refused rather than guessed at, as is one naming a
slot that does not exist.
Checked where it matters: the HOST formats a disk with a boot area and the
MACHINE fills it, then the reserved blocks are compared against zero. The
machine's allocator is the one that had no idea any of this was happening,
which is what makes that the check worth having. Six host checks besides,
including both halves of the superblock disagreeing.
Four things SplitLint wanted, and they build on each other.
EVERY WARNING NAMES ITS RULE, in brackets at the end the way a compiler
names the flag that produced it. Twelve rules, listed by --help. That makes
the other three possible: suppressions can name one rule and leave the line
honest about the others, the harness can assert on a rule's identity rather
than on the wording of its message, and --machine can print one tab
separated line per warning - file, line, rule, message, help - so nothing
downstream reads prose. This file's own output was parsed with regular
expressions three times in one day before it had a shape to rely on.
A CLEAN RUN SAYS SO:
No style warnings: 121 files checked against 12 rules.
It used to exit in silence, which does not tell you it found nothing - it
tells you nothing at all, and from outside the two are identical.
A MARKER THAT SILENCES NOTHING IS ITSELF REPORTED, as dead-suppression. An
exception that outlived whatever made it necessary is the thing the
required reason exists to prevent, and naming the wrong rule now gets you
both the warning you meant to silence and a note that your suppression is
doing nothing.
AND THE CORPUS IS HELD TO A BASELINE. Sixty one warnings are left in it
deliberately and nothing stopped a sixty second. Tests/lint-baseline.txt
records how many of each rule each file should produce, so a new one fails
make test while the sixty one stay quiet; confirmed by adding an INIA 0d0
to Say.asm and watching it name the file, the rule and the count. It counts
per file and rule rather than recording line numbers, because line numbers
would churn the whole baseline whenever anything was inserted above a
warning - the same reason cycle counts are stripped from recorded output
here. ./Tests/lint.sh --bless records it again.
One thing to know for next time: the rule name was inserted before the line
number at all twenty one call sites, and the signature was changed to match
rather than the twenty one call sites being fixed. (path, rule, line) reads
no worse than (path, line, rule) and one edit has fewer ways to go wrong
than twenty one.
SplitLint knew that CALL restores A, B and Data Pointers 0 to 2, so a
pointer set before a call is still set after it. That is true, and it made
the tool give advice that was correct today and unsafe to take.
Of the 178 redundant SETDs it found across the corpus, 122 were redundant
ONLY because of that restore - the shape is everywhere, because it is how a
helper is given its arguments:
SETD.0 SbfsBlock
SETD.2 SbfsFileStart
CALL sbfsSetWord
SETD.0 SbfsBlock <- flagged
Removing that last line is right until sbfsSetWord is reached with RCAL,
which restores nothing - and that is not hypothetical, it is what RCAL was
added to this machine for, measured at close to halving the assembler's
memory traffic. The failure would also be silent from the linter's side: it
forgets everything across an RCAL, so it would stop reporting while the
removals stayed removed.
So a claim now ends at any call, for pointers and for registers, the way a
claim about carry already did. 257 warnings become 127, and the redundant
SETDs 178 become 54 - which is exactly the number an independent count of
"no CALL in between" had arrived at separately.
The fixture gained a SETD and an INIA repeated across a CALL, which must
stay quiet, and the harness fails with the old behaviour put back. Two
mistakes worth recording: the new expectations first pointed at the LABEL
above the repeats rather than the repeats, which passes for free because
nothing ever warns about a label; and the block landed in the middle of
another check's comment, leaving that comment describing the code below it
instead of its own.
SplitLint reports valid assembly that has a shorter direct expression:
zero loads that could be RSTA or RSTB, Q moved through the stack where MVQA
would do, self-cancelling push and pop pairs, assignments overwritten
before use, unreachable fallthrough, one-byte pointer moves that could be
INCD or DECD, a branch to the label directly below it, a SETD reloading an
address the pointer already holds, and branches whose carry is known.
Its model is deliberately local and conservative: every label and every
directive forgets all known state, so a claim only ever lives inside a
straight-line region. It knows the calling convention - CALL forgets DP3
and keeps the rest, RCAL and SWI forget everything - and it shares
assembly.o with the assembler, so an added opcode cannot leave it holding a
private copy of the instruction table.
260 warnings across the corpus, of which three were wrong in the way that
matters: branchTest.asm and interruptFlagTest.asm exist to check that a
branch whose carry is known behaves correctly, so a diagnostic saying the
outcome is known is exactly right and exactly unwanted.
A line whose comment says "splitlint: <reason>" is now not reported on. THE
REASON IS REQUIRED and a bare marker is refused, because a suppression
nobody explained outlives whatever made it necessary. Suppressed warnings
are not counted, so --fatal-warnings does not fail on one, and the number
of them is printed at the end so the claim is visible rather than silent.
Tests/lint.sh checked a TOTAL: twenty three warnings expected, twenty three
found. That number stays right while the thing behind it goes wrong - a
rule that stopped firing while another fired twice would pass, and so would
a rule reporting at the wrong line. It now checks which warning came out
and at which line, that nothing else came out, and that the four lines
meant to stay quiet did. Confirmed by breaking one rule's message and
watching it name that rule: the old assertion passed the same sabotage,
because the warning still fired and the count never moved.
Written with the user while I was away; my part is the suppression
mechanism, the harness rewrite, and the three marks in the test programs.
HALT is terminal - stepCPU returns at once when the Halt Flag is up, so a
halted machine does not execute, service devices, or take an interrupt -
and that has to stay true, because every test ends with a halt and "halted"
is how a program says it has finished. The consequence was that SplitBit
had no way to wait at all. Every wait was a spin, and a spin is bus
traffic: 11.5% of Type over a 14K file on a disk of ten thousand cycles,
after read-ahead had already hidden three quarters of the latency.
WAIT is 0xFE, one byte, no operands, sitting under HALT where the
instruction that almost stops the machine belongs. Three decisions in it:
- A line already standing means there is nothing to wait for, so WAIT does
nothing. That is what makes test-then-wait race-free.
- Any line ends the wait, masked or not, so a program can sleep on a device
it has no handler for and read its status afterwards. Masking says who
answers a request, not whether it happened.
- A line that wakes the CPU without being dispatched is taken down by the
WAIT. Left standing it would be found by the next WAIT, which would
return at once - the program would spin exactly as before while looking
as though it slept.
Waiting is NOT a Status bit, and that is the trap avoided rather than a
gap: Status rides into the interrupt frame and comes back out, so a machine
interrupted mid-wait would return from its handler still waiting, and wait
again for what it had already been given. An internal field instead.
Idle cycles are counted apart from bus cycles and the halt line says so
when there are any, which is what makes the difference observable at all -
with the line-clearing removed the total moves by ONE cycle, 20,100 against
20,099, and only the idle half changes, halving to 9,976. A test on
totals could never have seen it. Tests/terminal.sh asks that question,
being the file for things a recorded output cannot see, and fails with the
clear removed while "both reads finished" still passes.
Three collisions, all found by building it:
- 0xFE was the assembler's "not an instruction" sentinel. getOpcode now
answers a negative NOT_AN_OPCODE, which is outside the range of every
possible answer instead of inside the unused part of it.
- 0xFE was also what faultTest and faultResumeTest executed to provoke a
fault. They now use 0xFD and say why, because they did not fail when it
became an instruction - they HUNG, having started sleeping instead.
- Keys.asm has had a label called "wait" for a year, and mnemonics are
matched uppercased. What that reported was "Branch without label" at the
BRQ thirty lines away. The assembler now refuses a label that is already
an instruction, at the label, by name; every instruction added takes a
word out of the space of label names, so this will happen again.
A parent is an entry index PLUS ONE in two bytes, so entry 65535 has no
parent number: adding one wraps to zero, and zero is the root. Eight
entries to a block, so 8192 directory blocks reaches it and SplitDisk
formatted that happily.
It does not fail by refusing, which is why it was worth chasing rather than
reasoning about. Reproduced on a disk built for it: mkdir /deep/child, with
/deep at entry 65535, printed 'Made "/deep/child" as entry 0' and put child
in the ROOT. Listing /deep then showed nothing, because the search is for a
parent of 65536 and the entry carries zero - so the same mkdir succeeded
again, and again, and five entries called /child piled up in the root.
Duplicate names in one directory are the one thing rename refuses outright,
on the grounds that a search answers with whichever it meets first and the
rest can never be reached; this manufactured them one per attempt.
8191 blocks is the most, giving 65528 entries. Refused when formatting and
again when reading, in both implementations, because a disk claiming more
was made by something that never checked. On the machine only the high byte
of the count has to be looked at: anything from 0x20 up is too many.
Three checks, all of which fail with their guard removed. The machine's
disk claims the size rather than having it, so the test image is 64 blocks
that lie rather than sixteen megabytes that do not - mounting is refused at
the geometry, which is read out of block 0.
Saving something that already exists writes a temporary, deletes the
original and gives the temporary its name, so that nothing is lost if the
writing fails. The temporary was told apart from a real file by being
called sbfs.part or sbfs.out - and those are legal names. Starting a save
deleted whatever answered to one as stale scratch, so saving anything at
all in a directory destroyed your own file of that name there, silently.
Flag bit 0x04 now says it. The property is not in the contents - the same
bytes become the finished file the instant the rename lands - so it belongs
in the entry, which is the thing the commit changes. sbfsCreateTempAt is
the door temporaries come in by, the commit writes the flags flat along
with the name, and cleanup wipes what it finds only if the entry says it is
ours. Anything else stops the save instead.
The bit is also the recovery. Both listings show an unfinished write rather
than sizing it, because the size in the entry is the room that was asked
for and not what was written: "<unfinished>" from dir, and a line from
SplitDisk saying the blocks are held and a rename brings the data back.
That was the gap in what the last commit documented - the data survived a
crash and nothing would show you where it was.
Four new agreement checks, three of which fail with the guards removed. The
fourth needed rebuilding first: both tests started on one disk, and the
first save ate the sbfs.part that was the second test's SOURCE, so the copy
failed for want of a file, never opened a stream, and passed while
reporting on nothing. A disk each. The fifth check forges the wreckage by
setting the flag on a finished file, since nothing here can crash a save
half way through.
No version bump: a committed file never carries the bit, so a disk this
writes is byte for byte the disk the old code wrote, which the whole-image
comparisons confirm. Only the wreckage differs, and older code reads that
as an ordinary file - which is what it did before.
The disk's status has always had a bit meaning "still going", and the header
beside it has always said to honour it. Nothing did, because nothing could: the
host finished the transfer inside the instruction that asked for it, so the bit
could never be seen up and asking about it was asking about something that
cannot happen.
--disk-cycles gives it a latency. The command is still checked at once, because
a refusal is not work - a block that is not there fails before any head moves -
but the transfer is remembered and done when the machine has run that far. Until
then the buffer holds the block BEFORE this one.
That last part is the point. A program that does not wait gets the wrong bytes
rather than an error, which is the failure the bit exists to prevent and the one
that would never have shown up. With a latency of two thousand, CosmOS could not
even mount: sbfsMount reads block zero and looks straight at the buffer.
deviceTick is the general shape rather than a disk feature. Called once per
instruction with the machine's clock, it lets anything whose moment has come
finish - which is what a display that refreshes, or a port that waits on the
host, would want in exactly the same way.
The filesystem watches the bit now, in one small routine reached with RCAL. That
is not decoration: what it hands back is the settled status in A, and CALL puts A
back the way it found it, so an ordinary call cannot carry the one thing this
exists to carry. Two bytes of Stack rather than ten, in a routine that runs on
every block the machine ever touches - the first place in the system where the
new call is the right one rather than merely a cheaper one.
The manifest takes a @N after a disk, the way it already takes :ro, so a test can
ask for a slow one. cosmosSlowDisk lists a directory at two thousand cycles a
block and gets the same listing as everything else, which is the whole assertion:
a filesystem that did not wait would print nonsense rather than fail.
Zero is the default and every other test runs at it. What waiting costs, on a
directory heavy run: 229k cycles at zero, 275k at five hundred, 415k at two
thousand, 1.16M at ten thousand.
Co-Authored-By: Claude Opus 5 <noreply@anthropic.com>
Claude-Session: https://claude.ai/code/session_01E2JrLzFvuFX9fgi1LDRjrW
A blit cost ten cycles, which were the five port writes that set it up. The
quarter of a kilobyte that moved cost nothing, and no hardware moves a quarter
of a kilobyte for nothing.
BANKS ARE SEPARATE MEMORIES, AND THAT IS WHAT SETS THE RATE. A move between two
of them can overlap its read and its write - fetch the next byte while the last
one is stored - so it settles at a byte a cycle. A move within one bank cannot,
and costs two. A fill has nothing to read and costs one whatever the banks are.
The odd cycle on each is the pipeline filling.
That is not a modelling choice so much as a reading of the structure the machine
already has: a Program to Data blit is inherently twice the rate of a Data to
Data one, and it is legible why.
Measured: 256 bytes is 297 cycles across banks and 518 within one, both
including the instructions that ask for it.
WHAT IT TAUGHT, which was not what I expected. Charging for movement costs the
native assembler 0.4 per cent and costs directory work 13.4. The assembler reads
a block and then thinks about it for a long time, so the move is amortised into
nothing; the filesystem reads a block in order to look at it and does nothing
else in between.
So the case for a blitter that runs alongside the CPU is weaker than it sounds.
Concurrency pays when there is other work to do during the transfer, and the
place that spends its time moving memory is exactly the place with nothing else
to do - it blits a block precisely so that it can read it. What that workload
wants is a FASTER controller, not a concurrent one: a wider data path halves the
wait, and the machine is waiting either way.
Video is the case that would still want concurrency, since a frame can be moved
while the next one is worked out. That is an argument about software nobody has
written yet, and it is now an argument with numbers on the other side of it.
The byte at a time port is charged too, for the byte it moves beyond reaching
the port. Nothing polls CTRL_STATUS, so the transfer stalls whoever asked for
it, which is the conservative reading and the one the software already assumes.
Co-Authored-By: Claude Opus 5 <noreply@anthropic.com>
Claude-Session: https://claude.ai/code/session_01E2JrLzFvuFX9fgi1LDRjrW
cycleCount used to tick once per instruction, so RSTA cost what SETD cost and a
CALL moving ten bytes of Stack cost what a branch cost. No machine anybody could
build works that way, and the emulator's job is to be the thing the hardware is
designed against.
Every touch of memory now goes through one of four accessors that charge for it:
fetching an opcode, fetching the bytes after it, reading or writing Data Memory,
and reaching a device port. One access, one cycle, nothing overlapped. The
accessors exist so the cost is counted where the access happens rather than in a
table of per instruction costs kept somewhere else - a table like that is a
second copy of what the code does, and the two drift.
The run loop spends a budget of cycles instead of running a count of
instructions, so the emulated rate means something: an instruction costs what it
touches, and a batch ends when the cycles are gone.
What the numbers say now: RSTA 1 and SETD 4, being one byte and four. LDA 3,
DPUA 2. CALL and RET together 24, RCAL and RRET together 8, because the first
pair moves twenty bytes of Stack and the second moves four. The average SplitBit
instruction costs 3.72 of these, measured over the native assembler assembling a
program.
And the measurement that prompted all of this: converting the filesystem's
hottest leaf routine to RCAL is 3.1 per cent cheaper on a directory heavy
workload. The old model said 0.0, which is what a model that cannot see memory
traffic must say about a change that is nothing else.
Three tests moved. settle() strips the cycle count from recorded output, so
nothing should have churned - but it was anchored to the start of a line and
replCalculator's last output has no newline on it, which leaves the halt message
mid line where the pattern never reached. Not anchored any more.
The two Life programs are bounded by a cycle count because they never end, and
that number was rescaled from 3,000,000 to 11,200,000 - the same amount of work
at 3.72 cycles to the instruction. Nothing about either program changed. No limit
reproduces the old output exactly, because the cut now lands elsewhere in a
frame, so they are recorded again rather than tuned to match.
Whether hardware overlaps a fetch with the end of the previous instruction is
left open on purpose. This is the conservative model; pipelining is a decision to
make while drawing the hardware, not one to inherit from an emulator.
Co-Authored-By: Claude Opus 5 <noreply@anthropic.com>
Claude-Session: https://claude.ai/code/session_01E2JrLzFvuFX9fgi1LDRjrW
-S writes every label and the address it was given, in address order.
Nothing else knows that. A program on the disk is bytes; the monitor can
disassemble it but has no idea what any of it is called. So counting which
addresses a program calls says a great deal and names nothing - the answer
arrives as a column of numbers and somebody works out by hand which routine
each one is inside.
It was deferred when the native assembler was planned, as a listing and symbol
dump nobody needed yet. Finding out where the assembler spends its time is what
needed it: the top six call targets were addresses until this existed and are
numStep, numCompare, tokGet, srcNext, numAddByte and clsSameName with it.
Sorted by address rather than by name, because the question asked of it is
always "what is at this address".
Co-Authored-By: Claude Opus 5 <noreply@anthropic.com>
Claude-Session: https://claude.ai/code/session_01E2JrLzFvuFX9fgi1LDRjrW
The six settled back on the twenty fourth, built now.
RCAL and RRET are a call that puts nothing back. CALL restores A, B and Data
Pointers 0 through 2, which costs ten bytes of Stack and is why a subroutine
here can only hand anything back through Q, DP3 or memory. RCAL costs two and
restores nothing, which is what a short leaf routine wants and is unsafe in
exactly the way the name says.
They are a pair because the frames are different sizes: returning from one
through the other walks the Stack to somewhere that was never a return address.
That was the user's correction to the original proposal, which had a raw call
and no raw return.
DPUA and DPDA offset a Data Pointer by A; DPUW and DPDW by A and B together,
most significant first. DPUP and DPDN take a byte written into the program, so
moving a pointer by something just worked out meant storing it and loading it
back. Down as well as up on symmetry grounds, which was also the user's call -
the argument against it came from counting uses in a corpus written under the
constraint.
The opcodes sit where they belong: 0x16 and 0x1E immediately below CALL and RET,
and 0x4E through 0x51 at the end of the Data Pointer family. All six fit shapes
that already existed, so instructiontable.py needed only set membership and both
machine side copies of the table regenerated from it unchanged.
Checked at every level it exists at: the emulator runs them, the host assembler
encodes them, the monitor disassembles all six with the right lengths, and the
assembler that runs on the machine builds a program using them byte for byte
identically to the host - and that program runs.
The recorded test measures what the two calls COST as well as what they put
back, because an RCAL that quietly did what CALL does would still return to the
right place. It does not survive that: returned through RRET, it hangs.
docs.sh can read a two word number now. The count of instructions taking a Data
Pointer went past twenty, and the pattern only allowed one word, so the check
would have reported that the manual had stopped saying it rather than that the
number was wrong.
Co-Authored-By: Claude Opus 5 <noreply@anthropic.com>
Claude-Session: https://claude.ai/code/session_01E2JrLzFvuFX9fgi1LDRjrW
SBFS version two adds directories out of space each entry had already set
aside: two of the four reserved bytes become a parent, and one of the seven
spare flag bits says an entry is a directory. The entry is still thirty two
bytes, so it still divides two hundred and fifty six and still never straddles
a block, and nothing in the block layer knows anything happened.
A directory is an entry with no blocks. That is what keeps the flat array of
entries the whole allocation map, which is the property the format is built
on: with files laid down contiguously, every block is inside some entry's
range or it is not, and an entry with no range is in nobody's way. There is
still no allocation table to consult and none to keep right.
THE PARENT IS AN INDEX PLUS ONE, so zero means the root. A version one disk has
zeroes in those bytes, and "in the root" is exactly where every file on a flat
disk is - so a version one image is already a valid version two image, with
nothing to convert and no tool to convert it with.
A disk is at the lowest version that describes what is on it. format makes a
version one disk and mkdir is what raises it, so everything built here stays
readable by a reader that has never heard of a directory right up until it
really does have one. That is what lets this land before the machine knows
anything: the whole existing suite passes untouched.
The tool gains mkdir and rmdir, and list, put, get and delete take paths. list
also now reports entries used against entries available, because a disk has two
ceilings and the entry one is the one nobody notices until it bites.
rmdir refuses a directory with anything in it, and that is not politeness:
parents are entry indices, a freed index gets handed out again, and the
children of a removed directory would reappear inside whatever took its place.
Co-Authored-By: Claude Opus 5 <noreply@anthropic.com>
Claude-Session: https://claude.ai/code/session_01E2JrLzFvuFX9fgi1LDRjrW
It always wrote .bin, whatever it had built. So assembling a loadable
program without -o produced Say.bin containing SBEX - a boot image name on a
file the machine cannot boot, in a repository whose whole convention is that
a .bin is started from and a .sbx is loaded.
Successfully wrote SplitBit boot image to "hello.bin".
Successfully wrote SplitBit loadable program to "Say.sbx".
programIsLoadable() already existed and is already what decides which
writer runs; the name now asks it too. Nothing in the build depended on the
old behaviour, because everything that assembles anything passes -o.
THE ASSEMBLER THAT RUNS ON SPLITBIT ALREADY DID IT THIS WAY. Two assemblers
naming their output differently from the same source is exactly the kind of
difference that wastes an afternoon, and the newer one was right.
Co-Authored-By: Claude Opus 5 <noreply@anthropic.com>
Claude-Session: https://claude.ai/code/session_01E2JrLzFvuFX9fgi1LDRjrW
"Binary" was doing three jobs. It meant an SPBT file that the machine starts
from; it meant whatever the assembler happened to produce, which is now
either that or a loadable program; and it meant a compiled host tool. A word
that means three things means none of them, and the first of the three has a
name already - this project has been calling them boot images for a while
and the manuals had not caught up.
Where it means an SPBT file -> boot image
Where it means either output -> output
Where it means a host executable -> left alone
Where it means base two -> left alone
The user facing messages move with it:
Error: No boot image specified.
Usage: ./SplitBit [OPTIONS] <boot image>
Error: This is not a SplitBit boot image.
Error: This boot image is in format version 2, and this emulator reads 1.
Successfully wrote SplitBit boot image to "hello.bin".
The assembler's own help was the interesting case. Its -o writes either
format, so "the binary" there was never right - it is "the output" now, and
the message that names the format is the one that says which it wrote.
No recorded output contained the word, so nothing needed re-blessing.
Checked before starting rather than after.
Co-Authored-By: Claude Opus 5 <noreply@anthropic.com>
Claude-Session: https://claude.ai/code/session_01E2JrLzFvuFX9fgi1LDRjrW
The quotes are gone by the time a token is classified, so checkIfKeyword's
test of token[0] == '#' matched the STRING "#Program" sitting in a program's
Data Segment. It was read as the directive: the segment silently changed in
the middle of the data, the string's nine bytes were charged to the Program
cursor instead of the Data one, and every label defined after it came out
nine bytes wrong - in a file that still had a valid header, a plausible
length, and nothing to say about any of it. The only symptom was a program
that jumped into the middle of an instruction.
This is the FOURTH of the family. A string spelling a mnemonic assembled as
that instruction; a string beginning with a zero was rejected as a malformed
literal; a string in the Program Segment was discarded in silence. The
instruction check and the literal check both carry a "not a STRING" guard
already. This one did not, so it has one now, and it lives inside
checkIfKeyword rather than at the call site so it cannot be left off again.
Nothing had ever triggered it, because nothing had ever needed a directive's
name as data. An assembler written FOR this machine necessarily does: it has
to compare tokens against "#Program" and "#Data". It was found by building
one and watching it fault on its second instruction.
Test stringKeyword puts every directive name in a program's data and prints
a label defined after them. Verified that it bites: without the fix the
assembler refuses the file outright.
Co-Authored-By: Claude Opus 5 <noreply@anthropic.com>
Claude-Session: https://claude.ai/code/session_01E2JrLzFvuFX9fgi1LDRjrW
A token is classified after its quotes have been stripped, so a string
literal reading "ADD" looked exactly like the ADD instruction and was
assembled as one. It failed with "attempting to assemble outside the
Program Segment", a message about a mistake nobody had made.
The literal and label checks were already guarded against strings and the
instruction check was not. Mnemonics match without regard to case, so
"or" and "and" were caught by this too, and those are ordinary enough
words to want in a message.
Third of its family, after a string beginning with '0' being read as a
malformed number and a string in the Program Segment being silently
discarded. All three have the same root.
Co-Authored-By: Claude Opus 5 <noreply@anthropic.com>
The loadable program format's header has been in the working tree since
the loader was written and has never been in the repository: the gitignore
rules matched Source/Assembler as a directory, so it was silently
untracked. Anything cloning this repository could not build it, because
secondPass.c includes this file.
This is the failure the previous commit describes, having already happened
once without being noticed.
Co-Authored-By: Claude Opus 5 <noreply@anthropic.com>
Overflowing/Underflowing A or B with increment or decrement instructions now sets the Carry Flag.
An increment or decrement of A or B that doesn't result in an overflow/underflow now clears the Carry Flag.
Q is no longer saved and restored in subroutine calls, allowing it to be used to return a one byte result.
Documentation updated with more detail.
Added DPUP and DPDN, which take a one byte operand specifying how far up or down to offset the Data Pointer.
Three Fibonacci generators using the print.asm library.
- 8 bit values printing in decimal representation.
- 16 bit values printing in hexadecimal representation.
- 32 bit values printing in hexadecimal representation.
Anachronaut
Anachronaut