15 Sep 2015
Over the weekend of September 5 through September 7, I had the rare opportunity to relax and write code for the Kestrel. I rather expected to continue to refine the ROM-based Forth image. Instead, I ended up writing a compiler for a simple programming language loosely inspired by previous generations of Machine Forth. I call it BSPL, short for Bootstrapping Systems Programming Language.
As I’ve written in the past, the RISC-V architecture is singularly difficult to write a Machine Forth dialect for. The 80386, 6809, 65816, 68000, and similar CISCy CPU families all provide sufficient built-in functionality that if you restrict your instruction choices enough, and in some cases make clever register allocations, you can treat these CPUs more or less as a stack architecture CPU. It’s quite rare for any Machine Forth primitive to expand to more than two CISC instructions, making a Machine Forth implementation for CISC CPUs easy. Software may not be the fastest possible, but the compiler will be small, convenient to write, substantially easier to port than even the simplest of C compilers, and utterly trivial to debug. You could reasonably get a Machine Forth compiler for x86 running in about two days worth of effort. I know, because I’ve done it, albeit for the Kestrel-2.
PowerPC, RISC-V, MIPS, SPARC, and other kinds of RISCy CPUs, however, don’t fare nearly as well. These CPUs provide primitives that are far too atomic, and thus, to emulate a stack CPU, you literally need to translate each Machine Forth primitive into a sequence of RISC instructions. For example, simply adding two numbers on the data stack involves the following code on RISC-V architectures:
ld x16, 0(dsp) ld x17, 8(dsp) add x16, x16, x17 addi dsp, dsp, 8 sd x16, 0(dsp)
Ouch. If we cache the top of the data stack in a CPU register, say X16, then we save two instructions:
ld x17, 0(dsp) ; DSP points to 2nd top of stack now add x16, x16, x17 addi dsp, dsp, 8
Still, we’re talking a minimum of three instructions just to add two numbers and make sure our data stack pointer is properly aligned. To see why this sucks so much, let’s look at the same code for the 68000:
add.l (a0)+, d0
Along a somewhat related vein, I don’t think I need to explain the impact on code bloat this has. It should be fairly obvious from the above example! The larger the code, the more cache pressure you have, and thus, the more cycles spent fetching cache lines instead of running your code. The RISC-V code is six times larger than the 68000 code.
Clearly, the naïve “macro expansion” method of compiling a Machine Forth dialect will not work for RISCy CPUs. We need something more sophisticated, yet still small and easy enough for a curious person to understand.
To finally realize a reasonable Machine Forth-like compiler for the RISC-V architecture, I combined several ideas from various sources, including one of my very own: first, I redefined the run-time execution model by considering a third stack; second, I abandoned the 1-pass compiler model and switched to a multi-pass compiler approach; third, while all this is going on, I maintain the intermediate program representation in a distinct “compile buffer”; and fourth, I made the evaluation stack ephemeral across basic-block boundaries.
One of my great insights came when I pondered, what if we used a three-stack virtual machine instead of a two-stack virtual machine? I can keep Forth’s traditional data and return stacks in memory; but, instead of conflating CPU registers with the data stack, I just treat CPU registers as a dedicated evaluation stack that’s kept strictly internal to the CPU. The compiler keeps track of which registers are in use, which register is known to be the current top of the evaluation stack, and other matters of statically-knowable, run-time bookkeeping.
The idea is simple enough. If we push a literal onto the stack, we allocate a CPU register to hold that value, and load the constant into that register. If we use a stacked datum, we use the corresponding register as a source to some instruction, and then mark that register free for later re-use.
Let’s pretend we have the input sequence
d# 1 d# 2 xor.
The first time
it allocates a new register (let’s say it’s X16),
places X16 onto the compiler’s evaluation stack,
then emits code to load X16 with the value
The second time
it allocates a new register (this time X17),
places X17 onto the compiler’s evaluation stack,
then emits code to load X17 with the value
(Aside: In practice, it doesn’t work this way; as you’ll read about below, these three steps are factored out into different compiler passes. However, this is what happens conceptually. I’ll get to the compiler’s nano-pass design later.)
You’d think that the Forth environment would have
pushed onto a stack somewhere, but it doesn’t.
The Forth environment doesn’t exist at this point.
However, the compiler knows that we just did something with
X16 and X17 in that order.
So, the compiler’s evaluation stack holds references to those two registers,
also in that order.
What happens when we
xor the two values?
The compiler knows that
xor takes two parameters (it’s a primitive),
and so consults the compiler’s evaluation stack
to see which registers those parameters are in.
It finds X16 and X17, with X17 on top.
So, it generates a
xor x16, x16, x17 instruction as output,
and pops X17 off the compiler’s evaluation stack.
This now leaves X16 as the new top of stack,
freeing X17 up for re-use in a subsequent computation.
BSPL currently reserves X16 through X31 for use as the evaluation stack. However, I should point out that I’ve written a number of non-trivial programs with BSPL, including unit testing tools and SD/MMC card device drivers. I’ve yet to see more than three of these registers used in real-world code.
For straight-ahead code, with no branches or other basic block boundaries, this actually results in surprisingly good code. It’s not as good as a graph-coloring or see-the-future register allocator, of course. It’s wholly ignorant of superscalar execution opportunities, among other things, and duplicate loads occur reasonably frequently. Still, it’s a far cry better than that mess I illustrated above or in my previous article on the topic. In most cases, a single Machine Forth primitive corresponds to a single CPU instruction.
Consider the following Forth code naïvely ported to run in BSPL:
variable foo : blort 1 foo @ if drop else 2 then 3 ;
After you run
blort, what is the state of the evaluation stack?
Forget any specific values, how deep is that stack?
Clearly, the precise behavior will depend on the state of
If it’s zero, you’ll have three things on the stack (1 2 3);
otherwise, just one (3).
This means that the 3 may appear either in register X16 or in X18,
depending on the state of
This literally means that any code following the IF-statement
cannot know where to find the top of stack.
Traditional compilers use various control-flow analysis techniques to “back propegate” preferred registers up into earlier segments of the generated code, to minimize register/memory synchronization overhead. For example, if I may over-generalize, phi-functions in static single assignment more or less gives the compiler a back-end directive that says, “choose a register, and substitute it for all these terms above.”
For now, I work around these problems by just punting on the problem all-together. At each basic-block boundary, I simply assume the state of the evaluation stack is completely empty. It doesn’t matter how many items I push into registers; when execution crosses a basic boundary block, such as a label or a jump instruction, we throw away everything already on the evaluation stack, and start allocating registers starting with X16 all over again.
This means you must explicitly load and store values to the data or return stack if you want the data to persist across basic blocks. For example, if we want to preserve the ANSI Forth semantics for the above if-statement, we’d write it as follows:
dword foo : blort d# 1 >d foo @ if d> else d# 2 >d then d# 3 >d ;
Take a look at the consequent of the IF-statement above.
Note that I use
d> to pop the 1 off the data stack,
which places it into a CPU register.
But, because I cross a basic block immediately afterwards,
by the time
d# 3 >d is executed,
it’s value is completely lost,
leaving only the 3 on the data stack, as intended.
Otherwise, during the alternate,
we push 2 onto the data stack,
so that by the time we leave the IF-statement,
both 1 and 2 are on the stack as we expect them to be.
d# word is a parsing word which
takes a decimal literal and compiles it as a literal push.
Beyond that, it’s unrelated to the data-stack accessors.
h# exists for hexadecimal numbers as well.)
Because it takes a control flow change to invoke a colon definition, the compiler assumes the evaluation stack is empty at the start of all definitions. For the same reason, the compiler must assume the evaluation stack is empty upon returning from an invoked word. All parameters are to exist on the data stack. Put in terms compatible with other contemporary programming languages, word invokations, IF-blocks, and BEGIN-blocks (not shown) are statements, not expressions.
Compared to traditional or CISC-targeting Machine Forth, BSPL (currently) can be pretty noisy due in large part to the need to manually manage parameter persistence. Despite this, I find it’s no worse than the assembler I used for the Kestrel-2. Indeed, despite the totally different designs, BSPL appears to the user to be a spiritual descendent of the Kestrel-2 assembler.
Here’s some code that prints an ISO 8829-1 character dump to the Kestrel-3 emulator’s debugger console:
\ Print characters in 16x16 matrix. \ Dump character to debugger port. : emit d> h# 0E00000000000000 c! ; \ Print special characters or control codes. : cr d# 13 >d emit d# 10 >d emit ; : space d# 32 >d emit ; \ Print a character followed by a space. \ Print a dot if the character is not printable. : ch 0 d@ d# 32 - -if d> d# $2E >d then emit ; : c 0 d@ >d ch space 0 d@ d# 1 + 0 d! ; \ Print 16 characters per row, then emit a carriage return. : r c c c c c c c c c c c c c c c c cr ; \ Starting with character code 0, print 16 character rows, \ then deadlock. : boot d# 0 >d r r r r r r r r r r r r r r r r begin again ;
The next contribution to the compiler came after I realized that a 1-pass compiler is simply impossible when your host CPU cannot always inline a numeric literal. Thus, I needed a multi-pass compiler, if for no other reason than accomodating space for literals. It turns out to be really handy for dealing with efficient return address and global pointer register management as well.
Before I go further, I want to make sure I provide appropriate attribution for this next contribution to my compiler. This is one of those ideas that’s so obvious in retrospect, but for some reason, it’s totally escaped me until now.
This idea was presented to me from a fellow homebrewer and 6502-entheusiast, Hans Franke, who brought this technique to my attention. It derives from a simple 2-pass assembler written by Lois Nachtmann for the 6502-based Elektor Junior. This assembler was tiny: it fit in a meager 2KB of memory, yet allowed forward references and other nice features typically associated with larger assemblers. The computer and its source code is described in a set of books available online. I don’t have these books myself; my understanding of the technique is based on information provided by Hans, who does have primary sources available to him.
Words compile not to native machine language, as most normal Forth environments would. Instead, they compile to an intermediate representation. This representation is basically a symbolc MS1 instruction set. At this point, each instruction compiled takes the following form:
n-1 8 7 0 +---------------------+--------+ | //// parameter //// | opcode | +---------------------+--------+
The low byte specifies an opcode, while all remaining bits in the cell (n >= 32; in my case, n=64) provides an operand for that opcode, if it takes one. (If not, the space is reserved, but goes unused.)
Instructions are laid down in a typical 1-pass compiler approach common to virtually all Machine Forth compilers. Here are some of the primitives BSPL supports as of this writing:
|noop||Take no action for the shortest, non-zero length of time possible.|
|lit||Offset to constant||Push literal onto evaluation stack.|
|label||Label index||Declares a label at this point in the program.|
|@8||Signed fetch byte.|
|@16||Signed fetch half-word.|
|@32||Signed fetch word.|
|@64||Signed fetch double-word.|
|goto||Label index||Unconditional jump.|
|gotoz||Label index||Jump if TOS is zero.|
|call||String index||Call subroutine by name.|
|rfs||Return from subroutine. (aka EXIT)|
|>r||Push item to R-stack.|
|r>||Pop item from R-stack.|
|r@||Cell offset||Fetch from R-stack.|
|r!||Cell offset||Store into R-stack.|
|>d||Push item to D-stack.|
|d>||Pop item from D-stack.|
|d@||Cell offset||Fetch from D-stack.|
|d!||Cell offset||Store into D-stack.|
|dup||Duplicate top of eval stack.|
|over||Duplicate second top of eval stack.|
|drop||Drop top of eval stack.|
|nip||Drop second top of eval stack.|
|rot||Rotate top three eval stack registers.|
|swap||Swap top two eval stack registers.|
|gotonz||Label index||Jump if TOS is non-zero.|
|2*||Shift TOS left by 1 bit.|
|u2/||Shift TOS right by 1 bit, filling in with 0 bits.|
|2/||Shift TOS right by 1 bit, preserving sign.|
|gotoge||Label index||Jump if TOS is greater than or equal to zero.|
|gvpea||Byte offset||Push effective address relative to GVP register.|
|strlit||String index||Push address and length for referenced string.|
Things differ from your common Machine Forth instruction set though.
Notice there’s an instruction called
Control flow happens relative to labels,
the index of which is declared using the
Note that a label may be declared after it’s referenced
(see the definition of
which means we would need at least two compiler passes to resolve them.
(Currently, BSPL emits assembly source listings,
relying on a subsequent call to an assembler to perform the 2nd pass.
However, nothing prevents this from being done inside the compiler itself.)
Subroutines are referenced by name, not by address.
When you invoke
the compiler interns the string
foo in a buffer,
and creates a unique token for it.
call then references that token in its parameter field.
The compile buffer will have several pointers referencing it.
+------------------+--------------+---/ | numeric literals | ... code ... | +------------------+--------------+---/ ^ ^ ^ | | | +- buf fi -+ ni -+
buf is the compile buffer itself. Executing that words pushes its address on the stack.
fi is the variable that points to the first instruction in the code space.
ni is the variable that points to the next instruction location, should you decide to compile another.
for those compiler passes that use it,
another variable called
ci refers to the current instruction being considered by that pass.
At all times,
ni are indices into the
not actually pointers.
Effective addresses are always computed from them,
so I think of them in terms of pointers anyway.
As you assemble individual instructions,
they are placed at
ni is bumped to the next location.
If a literal is pushed,
the whole image from
ni is moved
to make room for the new value at
Obviously, in this case, we need to update both
for later passes to use.
This can go on forever as long as
ni doesn’t overflow the buffer.
Several compiler passes exist to incrementally refine the code into a representation closer to a native RISC-V instruction stream. Arguably, the first pass is translating the source into symbolic form. All subsequent passes exist to massage the highest-level intermediate representation discussed previously into a form where one virtual instruction closely maps to a single RISC-V instruction.
After initial translation, the following passes occur in sequence. Each pass aims to be very small, thus easily manageable. I happened upon the idea of many tiny passes (as distinct from a few big passes) from this paper on Nanopasses.
The RISC-V architecture provides two ways to load constants into registers. I illustrate them below:
addi x10, x0, n ; -2048 <= n < 2048
ld x11, offset(x5) ; for all n
Optimally, if we want to push a number between -2048 and 2047 inclusive,
then we can just emit a single
Any value outside of that range
requires we store the literal
outside the normal control flow of the program,
where we use normal memory accessors to load it into the desired register.
Right now, BSPL treats all numeric literals uniformly.
I store each literal at the beginning of the compiled procedure.
This lets me statically compute the literal’s offset,
since I always know where the start of the procedure resides in memory.
Unfortunately, the CPU doesn’t at run-time;
we need to use the
to retrieve the current program counter into a “global pointer” register.
We could introduce
auipc in front of every literal load operation,
but this would cut performance significantly.
Small literals are used to bump pointers and decrement counters inside loops,
so it makes sense to optimize out all unnecessary
auipc instructions early on.
Further, because any called subroutines may reset the global pointer,
for these subroutines might use literals themselves,
we need to also insert it in front of the first literal load after any subroutine calls.
The logic behind
the lgp.fs module
implements this intelligence.
lgp.fs pass completes,
certain literal load instructions will be prefixed
with a LGP virtual instruction, like so:
+-------------+--------+ | | LGP | <-- inserted after this pass +-------------+--------+ | byte offset | LIT | +-------------+--------+ | label index | LABEL | +-------------+--------+ | | LGP | <-- inserted after this pass +-------------+--------+ | byte offset | LIT | +-------------+--------+ | byte offset | LIT | <-- Note lack of LGP here! +-------------+--------+ | label index | GOTOZ | +-------------+--------+ | | |
The reason we insert
LGP after the label
should become clear when you consider that
the label exists for the purpose of control flow change.
It’s possible that control will flow to that point
from somewhere such that the global pointer will not be set correctly.
The compiler plays it safe by making sure
LGP is inserted accordingly.
Since I’m compiling for a 64-bit RISC-V architecture,
all numeric literals are stored as double-words,
even though the finished program representation will use single-words for instructions.
It would be fairly easy to adjust
to recognize constants that fit the indicated range,
and to emit a new kind of instruction that directly inlines the constant.
I just haven’t done it out of laziness and a desire for semantic correctness.
Get something working first,
make it smaller/faster/better later.
I actually list two passes inside
the prolog-epilog.fs module.
The prolog pass basically asks the question,
“Do I need to incur the overhead of preserving the return address register?”
It works by scanning through the compiled set of instructions,
looking for any subroutine calls.
If no subroutine calls are detected,
then no need exists to preserve the
ra (return address) register.
Otherwise, we insert a
prolog virtual instruction at
In later passes, this will be expanded into something that preserves
ra on the run-time returns tack.
The epilog pass only runs if the prolog pass inserted a
It walks through the compiled code,
epilog virtual instructions in front of any subroutine returns,
if they’re needed at all.
As you might expect,
epilog instructions will eventually expand to code that restores the return address from the return stack.
So far, our goal in running all these passes has been to get our internal program representation more or less coincident with the final assembly language output. That is, to the greatest extent possible, we should be able to map one virtual instruction to one RISC-V instruction.
So far, however, we have not addressed the issue of register allocation. The register assignment pass works to allocate enumeration stack registers and assigns them as appropriate to source and/or destination slots for individual instructions.
It works by maintaining a compile-time stack.
As each instruction receives its consideration,
the pass will either allocate or dispose of CPU registers
according to normal stack management practices.
We could just use a simple counter for this job
were it not for the availability of enumeration stack permutation instructions,
dup, and so forth.
For each instruction between
the pass dispatches to a specific handler routine for that class of instruction.
A simple, linear scan through the instructions proves sufficient.
Up to now, all virtual instructions had one parameter field.
Some instructions now need a plurality of parameters.
+ requires two registers, a source and a destination.
The source register will be removed from the compile-time stack,
since it’s corresponding slot in the evaluation stack will no longer be allocated.
Others require a register field and its existing parameter field.
For such instructions, the word
scoot exists to vacate enough bits to hold register information.
scoot is responsible for vacating enough bits to hold register allocations.
Note that not all instruction types need scooting.
The before and after results of the register allocation pass for a few kinds of virtual instructions follows (LGPs and artifacts from other passes elided for clarity):
Before After 63 8 7 0 63 32 31 24 23 16 15 8 7 0 +-------------+--------+ +-------------+--------+--------+--------+--------+ | label index | LABEL | | label index | LABEL | +-------------+--------+ +-------------+--------+--------+--------+--------+ | byte offset | LIT | | byte offset | ////// | ////// | X16 | LIT | +-------------+--------+ +-------------+--------+--------+--------+--------+ | | @64 | | /////////// | ////// | ////// | X16 | @64 | +-------------+--------+ +-------------+--------+--------+--------+--------+ | byte offset | LIT | | byte offset | ////// | ////// | X17 | LIT | +-------------+--------+ +-------------+--------+--------+--------+--------+ | | ADD | | /////////// | ////// | X17 | X16 | ADD | +-------------+--------+ +-------------+--------+--------+--------+--------+ | | DUP | | /////////// | ////// | ////// | ////// | DUP | +-------------+--------+ +-------------+--------+--------+--------+--------+ | byte offset | LIT | | byte offset | ////// | ////// | X17 | LIT | +-------------+--------+ +-------------+--------+--------+--------+--------+ | | !64 | | /////////// | ////// | X17 | X16 | !64 | +-------------+--------+ +-------------+--------+--------+--------+--------+ | label index | GOTOZ | | label index | ////// | ////// | X16 | GOTOZ | +-------------+--------+ +-------------+--------+--------+--------+--------+
As you can see, what once appeared as a more-or-less pure stack-architecture sequence of instructions now looks quite close to RISC-V assembly language. We can now map these virtual instructions onto individual RISC-V instructions with greater ease.
Before continuing with the last pass in the BSPL compiler, I need to point out that a big opportunity for optimization exists.
If we wanted to provide additional (peephole) optimizations,
such as coalescing multiple bit-shift instructions into a single instruction with a larger-than-1 count,
LIT ADD @ or
LIT ADD ! sequences into single instructions,
one of the times to do that would be right now.
It wouldn’t even be that hard to do,
and it’d still retain the desirable characteristic
where Forth code maps closely to generated machine language.
Again, I haven’t implemented any of this logic because I wanted to focus on getting something that produced correct (even if slow) code first. But, as time progresses, you can bet that I’ll be writing additional optimization passes for BSPL.
At this point, most of the instructions in the compile buffer corresponds to one CPU instruction. If any optimization passes were run, then the program in the compile buffer should be as close to optimal as stack-architecture compiler technology can bring it. So, the final pass through the compile buffer produces the assembly language listing. It walks through every instruction in the compiler buffer, and emitting its textual representation.
Since we know every dword of memory between
fi correspond to numeric literals,
we know to just emit these literals as
dword directives in the source listing.
ni, we use a jump table to invoke code
specific to the opcode.
In most cases,
that code’s handler is just a one-line print routine,
dumping the listing to
And that’s all it takes to convert something resembling a Machine Forth dialect into a RISC-V instruction stream. The code won’t run as fast as something like modern-day C output, but right now, that’s not my intent. I just want something that is:
BSPL currently takes up 804 lines of code, of which 350 lines are commentary, leaving 444 lines of actual productive code. I’m willing to wager that this is pretty small by any compiler’s standards, let alone one that is multi-pass and capable of rudimentary optimizations in the future.
BSPL did take me longer than two days to get working. Closer to several weeks, including the odd bug-fix in my RISC-V emulator. However, it was still easier to get running than any C compiler project I looked at, including BCC, TCC, TinyC, GCC, CLang, and many others. Maybe some day, I can move on to porting a C compiler, or somehow figure out how to properly use the GNU toolchain in a Kestrel-3 compatible way. (If you’re not targeting Linux or Raspberry Pi, it seems, you’re SOL.) Until then, I look forward to many productive hours of actual coding, in Forth.
Software engineer by day. Amateur computer engineer by night. Founded the Kestrel Computer Project as a proof-of-concept back in 2007, with the Kestrel-1 computer built around the 65816 CPU. Since then, he's evolved the design to use a simple stack-architecture CPU with the Kestrel-2, and is now in the process of refining the design once more with a 64-bit RISC-V compatible engine in the Kestrel-3.
Samuel is or was:
Samuel seeks inspirations in many things, but is particularly moved by those things which moved or enabled him as a child. These include all things Commodore, Amiga, Atari, and all those old Radio-Electronics magazines he used to read as a kid.
Today, he lives in the San Francisco Bay Area with his beautiful wife, Steph, and four cats; 13, 6.5, Tabitha, and Panther.