nctref/DOCUMENTATION.md

# Nectar Reference Compiler Source Documentation

If you know this compiler took since 2019 to get to its current state, you will correctly guess that I don't really know what I am doing. Compiler literature, and online discussion, is abstract to the point where it is not useful for real-world processors. As a result, much of what you see in the source is the result of a lot of experimentation. I'm sure better methods are available to do the things within.

Basically, the compiler works by progressively iterating through the AST, turning it into a more primitive form step by step. This is necessary because machine code itself is primitive, and instructions typically have only 1-3 operands. Thanks to both this, and Nectar itself being a highly low-level language, the need for any IRs disappear. On the other hand, making sure the AST is in a correct state between steps isn't easy, and is the prime source of bugs.

Currently the compiler is designed with only i386+ processors in mind. I intend to add support for i286- and other exotic processors, but I honestly don't see it happening ever, especially if this remains a solo project. More RISC architectures with regular register files will be easier to add support for.

## AST structure

Starting with a Nectar source file, the compiler begins with the two common passes: lexing and parsing. Parsing exploits Nectar's syntax quirks, and may jump back and forth multiple times to fully parse a source file. This is necessary to avoid having to forward declare items. At the end, parsing returns what is called an AST in the source, although formally speaking the term is incorrectly used.

An AST node *may not be shared* by multiple other nodes. Also, the internal Nectar AST does not have scaling for pointer arithmetic; all pointers behave as `u8*`.

Each basic block is called a "chunk", likely a term I took from Lua. Basic blocks may contain one another; the least deep one within a function is called the top-level chunk (very important). Top-level chunks may contain other top-level chunks, because user-defined functions are within the "global scope", which is considered a function in itself.

After a chunk is finished parsing, all local variables in its scope are added to the flattened variables list of the top-level chunk's `ASTChunk` structure. Names may conflict, but at this point they're no longer important. Also worth mentioning is that this flat list contains `VarTableEntry` structs, even though VarTables are now irrelevant.

There's enough types of passes to push us to have a generic way to invoke the visitor pattern on the AST. Because passes may do many different things to the AST, including modify it, the definition of a generic visitor is very broad. Most functionality is unused by each pass.

    void generic_visitor(AST **nptr, AST *stmt, AST *stmtPrev, AST *chu, AST *tlc, void *ud, void(*handler)(AST**, AST*, AST*, AST*, AST*, void*));

`*nptr` is the actual node that is currently being visited. It is behind an additional indirection, because the node may be replaced by another.

If the current node is within a statement (most are), `stmt` is equal to that statement. `stmtPrev` is the previous statement. This is necessary for patching in the linked list of statements within a chunk during modification passes. If there is no previous statement, then the head pointer of the singly-linked list must be patched through the `chu` node. The `tlc` is the top-level chunk, which may be equal to `chu`.

## Dumbification

Once the AST is parsed, we move to machine-specific passes (in this case, i386). The idea of turning the AST progressively primitive is called "dumbification" in the source. The most simple example would be the following:

    a = -b

which should become

    a = b
    a = -a

Because the `neg` instruction on x86 is single-operand. If targeting an arch like MIPS, this specific dumbification would not be used, because one can use the 3-operand `subu` with the zero register.

Another rule is to extract function arguments and place them into local variables, but *only* if they do not form an x86 operand (for example `5` is ok because `push 5` exists).

Dumbification must be repeated until there are no more changes. The dumbification part of the source is responsible for making sure the resulting AST is "trivially compilable" to the machine code. This is actually non-trivial, because what is trivially compilable depends on which registers are used in the end (a variable colored as `edi`, `esi` or `ebp` cannot be used for 8-bit stores/loads). These details are not taken into account by dumbification.

A common bug when writing a dumbification rule is ending up with one that is always successful. If this happens, the compiler could become stuck endlessly dumbifying, which is nonsense.

Pre-dumbification is a single-time pass that takes a top-level chunk, and inserts loads for the function arguments. Such unconditional instructions are not efficient, but they work.

Putting all of this together, here is an example of nctref's dumbification of the following Fibonacci implementation, as of writing. Here is the main source, and the compiler's debug output:

    fibonacci: u32(u32 n) -> {
        if(n <= 1) {
            return n;
        }
        return fibonacci(n - 1) + fibonacci(n - 2);
    };

    @unimp fibonacci: u32(u32) {
    n = *(@stack + 4);
    if((n <= 1)) {
    return n;
    }
    $dumb2 = n;
    $dumb2 = ($dumb2 - 1);
    $dumb0 = fibonacci($dumb2);
    $dumb3 = n;
    $dumb3 = ($dumb3 - 2);
    $dumb1 = fibonacci($dumb3);
    $dumb0 = ($dumb0 + $dumb1);
    return $dumb0;
    };

`@unimp` is anything unimplemented in the AST debug printer, but it should say `u32(u32)`. `@stack` is an internal variable that points to the beginning of the current stack frame.

## Use-def chain

I hate these things. If you don't want to use static single assignment form, this is one alternative. Another is def-use chains, but both are horribly underdocumented.

For each variable, its UD chain is a list of each usage in the AST, with the corresponding potential definition of the variable at that use. For each potential definition that exists at that point, there is one UD element in the chain. Users include optimizers and codegen. The UD-chains are continually regenerated when needed by using the UD visitor on the top-level chunk.

As always, it's not that fuckin simple. Imagine the following pseudocode:

    x = 0;
    loop {
        x = x + 1;
        y = 5;
    }

The UD-chain code knows nothing about loops. It only cares whether something comes before or after, so it'll assume y is not in conflict with x, and they'll end up in the same register. Because of this, the parser must insert a so-called "loop guard", which will turn the AST into the following:

    x = 0;
    loop {
        x = x + 1;
        y = 5;
    }
    x;

That's one problem, but there's another:

    x = 0;
    loop {
        do something with x
        x = x + 1;
    }

Despite appearing after in the source, `x = x + 1` is a potential definition for everything in `do something with x`! This means the UD-chain generator must go through loops twice -- once with the upper definitions, and once with definitions from within the loop.

## Coloring

At this point we have a very distorted kind of Nectar AST in our function. We've got basic blocks and other familiar things, but all variables are in a flat list. These variables are essentially the "virtual registers" you hear a lot about. Because x86 only has six general-purpose registers, we must assign each of these variables (VarTableEntry structures, abbr. VTE) to a physical machine register.

This problem is a large area of study in itself, but a common approach is to imagine it as a graph coloring problem, where vertices are VTEs, and edges connect conflicting VTEs that cannot have the same color. Said edges are determined using the UD-chains of both VTEs.

The actual coloring algorithm used is Welsh-Powell, which sorts the VTEs/vertices by degree before attempting greedy coloring.

If there's more colors than there are physical registers, then we have a conflict, and must spill. There are two ways to do so: spill2var and spill2stack. The former is necessary on boundaries where suddenly a specific register/color must be used (e.g. returning in `eax`). The latter transforms every use of a local variable (`ASTExprVar` where its VTE is of type `VARTABLEENTRY_VAR`) into the form `@stack + n`.

If spill2stack is used, then CG must fail once so that dumbification can be applied again.

## Pre-coloring

I skipped forward a bit. In reality, coloring assumes that all registers have equal importance, which is never true. A return value must be in `eax`, the remainder of division must be in `edx`, etc. In 64-bit, the index of an argument determines in which register it may end up.

The pre-coloring visitor applies said rules to the AST, setting the colors in the VTE. It is completely plausible that a conflict can occur here, too, from two variables having overlapping live ranges and the same color, but it can also be from demanding more than one color from the same variable. In the latter case, the pre-coloring visitor gives up as soon as its detected. In both cases we do spill2var, not spill2stack, because spilling to the stack doesn't solve the pre-coloring problem.

## Callee-saved pass

If a function uses a callee-saved register, these must be stored and loaded at the correct times. This is done by modifying the AST in a special pass.

Of the four currently used registers, only `ebx` is callee-saved. A random variable assigned to `ebx` is chosen, and is saved to/loaded from the stack. The rule is written such that dumbification isn't necessary, unlike spill2stack.

## Code generation

FINALLY. This pass doesn't use `generic_visitor`, because it may consume multiple sibling AST nodes for emitting code. At this point there's nothing arcane or obscure; the code is pretty self-explanatory.

Using the same Fibonacci example as above, this is the result.

    global fibonacci
    fibonacci:
    mov edx, [esp + 4]
    cmp edx, 1
    ja .L0
    mov eax, edx
    ret
    .L0:
    mov eax, edx
    dec eax
    push ecx
    push edx
    push eax
    call fibonacci
    add esp, 4
    pop edx
    pop ecx
    mov ecx, eax
    mov eax, edx
    sub eax, 2
    push ecx
    push edx
    push eax
    call fibonacci
    add esp, 4
    pop edx
    pop ecx
    add ecx, eax
    mov eax, ecx
    ret

## Adding a Feature

When adding a feature, first write it out in Nectar in the ideal dumbified form. Make sure this compiles correctly. Afterward, implement dumbification rules so that code can be written in any fashion. If specific colorings are required, then the pre-coloring and spill2var passes must be updated. The following is an example with multiplication, as this is what I'm adding as of writing.

Note the way `mul` works on x86. Firstly, one of the operands is the destination, because `mul` is a 2-op instruction. Secondly, the other operand may not be an immediate, because the operand is defined as r/m (register or memory), so if the second operand is a constant, it must be spilled into a variable (`varify` in `dumberdowner.c`). Thirdly, this destination must be the A register, so one of the operands must be pre-colored to A. Fourthly, `mul` clobbers the D register with the high half of the product. In other words, we have an instruction with *two* output registers, which the Nectar AST does not support. But we can't have the register allocator assign anything to D here.

To account for this, we can have a second assignment statement right next to the multiplication. Because the main multiplication clobbers the source operand, the mulhi assignment must come before the main mul. Putting all this together, this is the canonical way to do `z = x * y` with an x86 target:

    z = x;
    w = z *^ y;
    z = z * y;

Now we must modify the pre-coloring pass to make sure `z` is marked as A and `w` as D. In case such pre-coloring is impossible, the `spill2var` pass must also be modified to spill whatever variables prevent this coloring into another variable (NOT into the stack). If this is the last use of `y`, then it is fine for `y` to be assigned to D.

Lastly, the codegen pass must recognize the sequence `w = z *^ y; z = z * y;` and emit a single `mul` instruction.

In `cg.c` is a function called `xop`, which returns an x86 operand string, given a trivially compilable Nectar expression. Because we've guaranteed the other operand may not be a constant, we do not need to check the XOP type, but it's a good idea to insert `assert`s and `abort`s everywhere to prevent hard-to-find bugs.

Once all that is done and tested, now we can add the following dumbification rules: all binary operations with the operand `AST_BINOP_MUL` or `AST_BINOP_MULHI` must be the whole expression within an assignment statement. If not, extract into a separate assignment & new variable with `varify`. The destination of the assignment, and both operands of the binary operation must be of type `AST_EXPR_VAR`, with their corresponding variables being of type `VARTABLEENTRY_VAR`, not `VARTABLEENTRY_SYMBOL` or `VARTABLEENTRY_TYPE`. If any of those don't apply, `varify` the offenders. Each such assignment have a neighboring, symmetric assignment, so that both A and D are caught by the pre-coloring pass.