---
title: Compiling a Functional Language Using C++, Part 12 - Let/In and Lambdas
date: 2020-06-21T00:50:07-07:00
tags: ["C and C++", "Functional Languages", "Compilers"]
description: "In this post, we extend our language with let/in expressions and lambda functions."
---
Now that our language's type system is more fleshed out and pleasant to use, it's time to shift our focus to the ergonomics of the language itself. I've been mentioning `let/in` and __lambda__ expressions for a while now. The former will let us create names for expressions that are limited to a certain scope (without having to create global variable bindings), while the latter will allow us to create functions without giving them any name at all.
Let's take a look at `let/in` expressions first, to make sure we're all on the same page about what it is we're trying to implement. Let's start with some rather basic examples, and then move on to more complex ones. A very basic use of a `let/in` expression is, in Haskell:
```Haskell
let x = 5 in x + x
```
In the above example, we bind the variable `x` to the value `5`, and then refer to `x` twice in the expression after the `in`. The whole snippet is one expression, evaluating to what the `in` part evaluates to. Additionally, the variable `x` does not escape the expression -
{{< sidenote "right" "used-note" "it cannot be used anywhere else." >}}
Unless, of course, you bind it elsewhere; naturally, using x here does not forbid you from re-using the variable.
{{< /sidenote >}}
Now, consider a slightly more complicated example:
```Haskell
let sum xs = foldl (+) 0 xs in sum [1,2,3]
```
Here, we're defining a _function_ `sum`,
{{< sidenote "right" "eta-note" "which takes a single argument:" >}}
Those who favor the
point-free
programming style may be slightly twitching right now, the words eta reduction swirling in their mind. What do you know, fold-based sum is even one of the examples on the Wikipedia page! I assure you, I left the code as you see it deliberately, to demonstrate a principle.
{{< /sidenote >}} the list to be summed. We will want this to be valid in our language, as well. We will soon see how this particular feature is related to lambda functions, and why I'm covering these two features in the same post.
Let's step up the difficulty a bit more, with an example that,
{{< sidenote "left" "translate-note" "though it does not immediately translate to our language," >}}
The part that doesn't translate well is the whole deal with patterns in function arguments, as well as the notion of having more than one equation for a single function, as is the case with safeTail.
It's not that these things are impossible to translate; it's just that translating them may be worthy of a post in and of itself, and would only serve to bloat and complicate this part. What can be implemented with pattern arguments can just as well be implemented using regular case expressions; I dare say most "big" functional languages actually just convert from the former to the latter as part of the compillation process.
{{< /sidenote >}} illustrates another important principle:
```Haskell {linenos=table}
let
safeTail [] = Nothing
safeTail [x] = Just x
safeTail (_:xs) = safeTail xs
myTail = safeTail [1,2,3,4]
in
myTail
```
The principle here is that definitions in `let/in` can be __recursive and polymorphic__. Remember the note in
[part 10]({{< relref "10_compiler_polymorphism.md" >}}) about
[let-polymorphism](https://en.wikipedia.org/wiki/Hindley%E2%80%93Milner_type_system#Let-polymorphism)? This is it: we're allowing polymorphic variable bindings, but only when they're bound in a `let/in` expression (or at the top level).
The principles demonstrated by the last two snippets mean that compiling `let/in` expressions, at least with the power we want to give them, will require the same kind of dependency analysis we had to go through when we implemented polymorphically typed functions. That is, we will need to analyze which functions calls which other functions, and typecheck the callees before the callers. We will continue to represent callee-caller relationships using a dependency graph, in which nodes represent functions, and an edge from one function node to another means that the former function calls the latter. Below is an image of one such graph:
{{< figure src="fig_graph.png" caption="Example dependency graph without `let/in` expressions." >}}
Since we want to typecheck callees first, we effectively want to traverse the graph in reverse
topological order. However, there's a slight issue: a topological order is only defined for acyclic graphs, and it is very possible for functions in our language to mutually call each other. To deal with this, we have to find groups of mutually recursive functions, and and treat them as a single unit, thereby eliminating cycles. In the above graph, there are two groups, as follows:
{{< figure src="fig_colored_ordered.png" caption="Previous depndency graph with mutually recursive groups highlighted." >}}
As seen in the second image, according to the reverse topological order of the given graph, we will typecheck the blue group containing three functions first, since the sole function in the orange group calls one of the blue functions.
Things are more complicated now that `let/in` expressions are able to introduce their own, polymorphic and recursive declarations. However, there is a single invariant we can establish: function definitions can only depend on functions defined at the same time as them. That is, for our purposes, functions declared in the global scope can only depend on other functions declared in the global scope, and functions declared in a `let/in` expression can only depend on other functions declared in that same expression. That's not to say that a function declared in a `let/in` block inside some function `f` can't call another globally declared function `g` - rather, we allow this, but treat the situation as though `f` depends on `g`. In contrast, it's not at all possible for a global function to depend on a local function, because bindings created in a `let/in` expression do not escape the expression itself. This invariant tells us that in the presence of nested function definitions, the situation looks like this:
{{< figure src="fig_subgraphs.png" caption="Previous depndency graph augmented with `let/in` subgraphs." >}}
In the above image, some of the original nodes in our graph now contain other, smaller graphs. Those subgraphs are the graphs created by function declarations in `let/in` expressions. Just like our top-level nodes, the nodes of these smaller graphs can depend on other nodes, and even form cycles. Within each subgraph, we will have to perform the same kind of cycle detection, resulting in something like this:
{{< figure src="fig_subgraphs_colored_all.png" caption="Augmented dependency graph with mutually recursive groups highlighted." >}}
When typechecking a function, we must be ready to perform dependency analysis at any point. What's more is that the free variable analysis we used to perform must now be extended to differentiate between free variables that refer to "nearby" definitions (i.e. within the same `let/in` expression), and "far away" definitions (i.e. outside of the `let/in` expression). And speaking of free variables...
What do we do about variables that are captured by a local definition? Consider the following snippet:
```Haskell {linenos=table}
addToAll n xs = map addSingle xs
where
addSingle x = n + x
```
In the code above, the variable `n`, bound on line 1, is used by `addSingle` on line 3. When a function refers to variables bound outside of itself (as `addSingle` does), it is said to be _capturing_ these variables, and the function is called a _closure_. Why does this matter? On the machine level, functions are represented as sequences of instructions, and there's a finite number of them (as there is finite space on the machine). But there is an infinite number of `addSingle` functions! When we write `addToAll 5 [1,2,3]`, `addSingle` becomes `5+x`. When, on the other hand, we write `addToAll 6 [1,2,3]`, `addSingle` becomes `6+x`. There are certain ways to work around this - we could, for instance, dynamically create machine code in memory, and then execute it (this is called [just-in-time compilation](https://en.wikipedia.org/wiki/Just-in-time_compilation)). This would end up with a collections of runtime-defined functions that can be represented as follows:
```Haskell {linenos=table}
-- Version of addSingle when n = 5
addSingle5 x = 5 + x
-- Version of addSingle when n = 6
addSingle6 x = 6 + x
-- ... and so on ...
```
But now, we end up creating several functions with almost identical bodies, with the exception of the free variables themselves. Wouldn't it be better to perform the well-known strategy of reducing code duplication by factoring out parameters, and leaving only one instance of the repeated code? We would end up with:
```Haskell {linenos=table}
addToAll n xs = map (addSingle n) xs
addSingle n x = n + x
```
Observe that we no longer have the "infinite" number of functions - the infinitude of possible behaviors is created via currying. Also note that `addSingle`
{{< sidenote "right" "global-note" "is now declared at the global scope," >}}
Wait a moment, didn't we just talk about nested polymorphic definitions, and how they change our typechecking model? If we transform our program into a bunch of global definitions, we don't need to make adjustments to our typechecking.
This is true, but why should we perform transformations on a malformed program? Typechecking before pulling functions to the global scope will help us save the work, and breaking down one dependency-searching problem (which is \(O(n^3)\) thanks to Warshall's) into smaller, independent problems may even lead to better performance. Furthermore, typechecking before program transformations will help us come up with more helpful error messages.
{{< /sidenote >}} and can be transformed into a sequence of instructions just like any other global function. It has been pulled from its `where` (which, by the way, is pretty much equivalent to a `let/in`) to the top level.
Now, see how `addSingle` became `(addSingle n)`? If we chose to rewrite the
program this way, we'd have to find-and-replace every instance of `addSingle`
in the function body, which would be tedious and require us to keep
track of shadowed variables and the like. Also, what if we used a local
definition twice in the original piece of code? How about something like this:
```Haskell {linenos=table}
fourthPower x = square * square
where
square = x * x
```
Applying the strategy we saw above, we get:
```Haskell {linenos=table}
fourthPower x = (square x) * (square x)
square x = x * x
```
This is valid, except that in our evaluation model, the two instances
of `(square x)` will be built independently of one another, and thus,
will not be shared. This, in turn, will mean that `square` will be called
twice, which is not what we would expect from looking at the original program.
This isn't good. Instead, why don't we keep the `where`, but modify it
as follows:
```Haskell {linenos=table}
fourthPower x = square * square
where square = square' x
square' x = x * x
```
This time, assuming we can properly implement `where`, the call to
`square' x` should only occur once. Though I've been using `where`,
which leads to less clutter in Haskell code, the exact same approach applies
to `let/in`, and that's what we'll be using in our language.
This technique of replacing captured variables with arguments, and pulling closures into the global scope to aid compilation, is called [Lambda Lifting](https://en.wikipedia.org/wiki/Lambda_lifting). Its name is no coincidence - lambda functions need to undergo the same kind of transformation as our nested definitions (unlike nested definitions, though, lambda functions need to be named). This is why they are included in this post together with `let/in`!
What are lambda functions, by the way? A lambda function is just a function
expression that doesn't have a name. For example, if we had Haskell code like
this:
```Haskell
double x = x + x
doubleList xs = map double xs
```
We could rewrite it using a lambda function as follows:
```Haskell
doubleList xs = map (\x -> x + x) xs
```
As you can see, a lambda is an expression in the form `\x -> y` where `x` can
be any variable and `y` can be any expression (including another lambda).
This represents a function that, when applied to a value `x`, will perform
the computation given by `y`. Lambdas are useful when creating single-use
functions that we don't want to make globally available.
Lifting lambda functions will effectively rewrite our program in the
opposite direction to the one shown, replacing the lambda with a reference
to a global declaration which will hold the function's body. Just like
with `let/in`, we will represent captured variables using arguments
and partial appliciation. For instance, when starting with:
```Haskell
addToAll n xs = map (\x -> n + x) xs
```
We would output the following:
```Haskell
addToAll n xs = map (lambda n) xs
lambda n x = n + x
```
### Implementation
Now that we understand what we have to do, it's time to jump straight into
doing it. First, we need to refactor our current code to allow for the changes
we're going to make; then, we will use the new tools we defined to implement `let/in` expressions and lambda functions.
#### Infrastructure Changes
When finding captured variables, the notion of _free variables_ once again
becomes important. Recall that a free variable in an expression is a variable
that is defined outside of that expression. Consider, for example, the
expression:
```Haskell
let x = 5 in x + y
```
In this expression, `x` is _not_ a free variable, since it's defined
in the `let/in` expression. On the other hand, `y` _is_ a free variable,
since it's not defined locally.
The algorithm that we used for computing free variables was rather biased.
Previously, we only cared about the difference between a local variable
(defined somewhere in a function's body, or referring to one of the function's
parameters) and a global variable (referring to a global function).
This shows in our code for `find_free`. Consider, for example, this snippet:
{{< codelines "C++" "compiler/11/ast.cpp" 33 36 >}}
We created bindings in our type environment whenever we saw a new variable
being introduced, which led us to only count variables that we did not bind
_anywhere_ as 'free'. This approach is no longer sufficient. Consider,
for example, the following Haskell code:
```Haskell {linenos=table}
someFunction x =
let
y = x + 5
in
x*y
```
We can see that the variable `x` is introduced on line 1.
Thus, our current algorithm will happily store `x` in an environment,
and not count it as free. But clearly, the definition of `y` on line 3
captures `x`! If we were to lift `y` into global scope, we would need
to pass `x` to it as an argument. To fix this, we have to separate the creation
and assignment of type environments from free variable detection. Why
don't we start with `ast` and its descendants? Our signatures become:
```C++
void ast::find_free(std::set& into);
type_ptr ast::typecheck(type_mgr& mgr, type_env_ptr& env);
```
For the most part, the code remains unchanged. We avoid
using `env` (and `this->env`), and default to marking
any variable as a free variable:
{{< codelines "C++" "compiler/12/ast.cpp" 39 41 >}}
Since we no longer use the environment, we resort to an
alternative method of removing bound variables. Here's
`ast_case::find_free`:
{{< codelines "C++" "compiler/12/ast.cpp" 169 181 >}}
For each branch, we find the free variables. However, we
want to avoid marking variables that were introduced through
pattern matching as free (they are not). Thus, we use `pattern::find_variables`
to see which of the variables were bound by that pattern,
and remove them from the list of free variables. We
can then safely add the list of free variables in the pattern to the overall
list of free variables. Other `ast` descendants experience largely
cosmetic changes (such as the removal of the `env` parameter).
Of course, we must implement `find_variables` for each of our `pattern`
subclasses. Here's what I got for `pattern_var`:
{{< codelines "C++" "compiler/12/ast.cpp" 402 404 >}}
And here's an equally terse implementation for `pattern_constr`:
{{< codelines "C++" "compiler/12/ast.cpp" 417 419 >}}
We also want to update `definition_defn` with this change. Our signatures
become:
```C++
void definition_defn::find_free();
void definition_defn::insert_types(type_mgr& mgr, type_env_ptr& env, visibility v);
```
We'll get to the `visiblity` parameter later. The implementations
are fairly simple. Just like `ast_case`, we want to erase each function's
parameters from its list of free variables:
{{< codelines "C++" "compiler/12/definition.cpp" 13 18 >}}
Since `find_free` no longer creates any type bindings or environments,
this functionality is shouldered by `insert_types`:
{{< codelines "C++" "compiler/12/definition.cpp" 20 32 >}}
Now that free variables are properly computed, we are able to move on
to bigger and better things.
#### Nested Definitions
At present, our code for typechecking the whole program is located in
`main.cpp`:
{{< codelines "C++" "compiler/11/main.cpp" 43 61 >}}
This piece of code goes on. We now want this to be more general. Soon, `let/in`
expressions with bring with them definitions that are inside other definitions,
which will not be reachable at the top level. The fundamental topological
sorting algorithm, though, will remain the same. We can abstract a series
of definitions that need to be ordered and then typechecked into a new struct,
`definition_group`:
{{< codelines "C++" "compiler/12/definition.hpp" 73 83 >}}
This will be exactly like a list of `defn`/`data` definitions we have at the
top level, except now, it can also occur in other places, like `let/in`
expressions. Once again, ignore for the moment the `visibility` field.
The way we defined function ordering requires some extra work from
`definition_group`. Recall that conceptually, functions can only depend
on other functions defined in the same `let/in` expression, or, more generally,
in the same `definition_group`. This means that we now classify free variables
in definitions into two categories: free variables that refer to "nearby"
definitions (i.e. definitions in the same group) and free variables that refer
to "far away" definitions. The "nearby" variables will be used to do
topological ordering, while the "far away" variables can be passed along
further up, perhaps into an enclosing `let/in` expression (for which "nearby"
variables aren't actually free, since they are bound in the `let`). We
implement this partitioning of variables in `definition_group::find_free`:
{{< codelines "C++" "compiler/12/definition.cpp" 94 105 >}}
Notice that we have added a new `nearby_variables` field to `definition_defn`.
This is used on line 101, and will be once again used in `definition_group::typecheck`. Speaking of `typecheck`, let's look at its definition:
{{< codelines "C++" "compiler/12/definition.cpp" 107 145 >}}
This function is a little long, but conceptually, each `for` loop
contains a step of the process:
* The first loop declares all data types, so that constructors can
be verified to properly reference them.
* The second loop creates all the data type constructors.
* The third loop adds edges to our dependency graph.
* The fourth loop performs typechecking on the now-ordered groups of mutually
recursive functions.
* The first inner loop inserts the types of all the functions into the environment.
* The second inner loop actually performs typechecking.
* The third inner loop makes as many things polymorphic as possible.
We can now adjust our `parser.y` to use a `definition_group` instead of
two global vectors. First, we declare a global `definition_group`:
{{< codelines "C++" "compiler/12/parser.y" 10 10 >}}
Then, we adjust `definitions` to create `definition_group`s:
{{< codelines "text" "compiler/12/parser.y" 59 68 >}}
We can now adjust `main.cpp` to use the global `definition_group`. Among
other changes (such as removing `extern` references to `vector`s, and updating
function signatures) we also update the `typecheck_program` function:
{{< codelines "C++" "compiler/12/main.cpp" 41 49 >}}
Now, our code is ready for typechecking nested definitions, but not for
compiling them. The main thing that we still have to address is the addition
of new definitions to the global scope. Let's take a look at that next.
#### Global Definitions
We want every function (and even non-function definitions that capture surrounding
variables), regardless of whether or not it was declared in the global scope,
to be processed and converted to LLVM code. The LLVM code conversion takes
several steps. First, the function's AST is translated into G-machine
instructions, which we covered in [part 5]({{< relref "05_compiler_execution.md" >}}),
by a process we covered in [part 6]({{< relref "06_compiler_compilation.md" >}}).
Then, an LLVM function is created for every function, and registered globally.
Finally, the G-machine instructions are converted into LLVM IR, which is
inserted into the previously created functions. These things
can't be done in a single pass: at the very least, we can't start translating
G-machine instructions into LLVM IR until functions are globally declared,
because we would otherwise have no means of referencing other functions. It
makes sense to me, then, to pull out all the 'global' definitions into
a single top-level list (perhaps somewhere in `main.cpp`).
Let's start implementing this with a new `global_scope` struct. This
struct will contain all of the global function and constructor definitions:
{{< codelines "C++" "compiler/12/global_scope.hpp" 42 55 >}}
This struct will allow us to keep track of all the global definitions,
emitting them as we go, and then coming back to them as necessary.
There are also signs of another piece of functionality: `occurence_count`
and `mangle_name`. These two will be used to handle duplicate names.
We cannot have two global functions named the same thing, but we can
easily imagine a situation in which two separate `let/in` expressions define
a variable like `x`, which then needs to be lifted to the global scope. We
resolve such conflicts by slightly changing - "mangling" - the name of
one of the resulting global definitions. We allow the first global definition
to be named the same as it was originally (in our example, this would be `x`).
However, if we detect that a global definition `x` already exists (we
track this using `occurence_count`), we rename it to `x_1`. Subsequent
global definitions will end up being named `x_2`, `x_3`, and so on.
Alright, let's take a look at `global_function` and `global_constructor`.
Here's the former:
{{< codelines "C++" "compiler/12/global_scope.hpp" 11 27 >}}
There's nothing really surprising here: all of the fields
are reminiscent of `definition_defn`, though some type-related variables
are missing. We also include the three compilation-related methods,
`compile`, `declare_llvm`, and `generate_llvm`, which were previously in `definition_defn`. Let's look at `global_constructor` now:
{{< codelines "C++" "compiler/12/global_scope.hpp" 29 40 >}}
This maps pretty closely to a single `definition_data::constructor`.
There's a difference here that is not clear at a glance, though. Whereas
the `name` in a `definition_defn` or `definition_data` refers to the
name as given by the user in the code, the `name` of a `global_function`
or `global_constructor` has gone through mangling, and thus, should be
unique.
Let's now look at the implementation of these structs' methods. The methods
`add_function` and `add_constructor` are pretty straightforward. Here's
the former:
{{< codelines "C++" "compiler/12/global_scope.cpp" 39 43 >}}
And here's the latter:
{{< codelines "C++" "compiler/12/global_scope.cpp" 45 49 >}}
In both of these functions, we return a reference to the new global
definition we created. This helps us access the mangled `name` field,
and, in the case of `global_function`, inspect the `ast_ptr` that represents
its body.
Next, we have `global_scope::compile` and `global_scope::generate_llvm`,
which encapsulate these operations on all global definitions. Their
implementations are very straightforward, and are similar to the
`gen_llvm` function we used to have in our `main.cpp`:
{{< codelines "C++" "compiler/12/global_scope.cpp" 51 67 >}}
Finally, we have `mangle`, which takes care of potentially duplicate
variable names:
{{< codelines "C++" "compiler/12/global_scope.cpp" 69 83 >}}
Let's move on to the global definition structs.
The `compile`, `declare_llvm`, and `generate_llvm` methods for
`global_function` are pretty much the same as those that we used to have
in `definition_defn`:
{{< codelines "C++" "compiler/12/global_scope.cpp" 4 24 >}}
The same is true for `global_constructor` and its method `generate_llvm`:
{{< codelines "C++" "compiler/12/global_scope.cpp" 26 37 >}}
Recall that in this case, we need not have two methods for declaring
and generating LLVM, since constructors don't reference other constructors,
and are always generated before any function definitions.
#### Visibility
Should we really be turning _all_ free variables in a function definition
into arguments? Consider the following piece of Haskell code:
```Haskell {linenos=table}
add x y = x + y
mul x y = x * y
something = mul (add 1 3) 3
```
In the definition of `something`, `mul` and `add` occur free.
A very naive lifting algorithm might be tempted to rewrite such a program
as follows:
```Haskell {linenos=table}
add x y = x + y
mul x y = x * y
something' add mul = mul (add 1 3) 3
something = something' add mul
```
But that's absurd! Not only are `add` and `mul` available globally,
but such a rewrite generates another definition with free variables,
which means we didn't really improve our program in any way. From this
example, we can see that we don't want to be turning reference to global
variables into function parameters. But how can we tell if a variable
we're trying to operate on is global or not? I propose a flag in our
`type_env`, which we'll augment to be used as a symbol table. To do
this, we update the implementation of `type_env` to map variables to
values of a struct `variable_data`:
{{< codelines "C++" "compiler/12/type_env.hpp" 14 23 >}}
The `visibility` enum is defined as follows:
{{< codelines "C++" "compiler/12/type_env.hpp" 11 11 >}}
As you can see from the above snippet, we also added a `mangled_name` field
to the new `variable_data` struct. We will be using this field shortly. We
also add a few methods to our `type_env`, and end up with the following:
{{< codelines "C++" "compiler/12/type_env.hpp" 32 45 >}}
We will come back to `find_free` and `find_free_except`, as well as
`set_mangled_name` and `get_mangled_name`. For now, we just adjust `bind` to
take a visibility parameter that defaults to `local`, and implement
`is_global`:
{{< codelines "C++" "compiler/12/type_env.cpp" 27 32 >}}
Remember the `visibility::global` in `parser.y`? This is where that comes in.
Specifically, we recall that `definition_defn::insert_types` is responsible
for placing function types into the environment, making them accessible
during typechecking later. At this time, we already need to know whether
or not the definitions are global or local (so that we can create the binding).
Thus, we add `visibility` as a parameter to `insert_types`:
{{< codelines "C++" "compiler/12/definition.hpp" 44 44 >}}
Since we are now moving from manually wrangling definitions towards using
`definition_group`, we make it so that the group itself provides this
argument. To do this, we add the `visibility` field from before to it,
and set it in the parser. One more thing: since constructors never
capture variables, we can always move them straight to the global
scope, and thus, we'll always mark them with `visibility::global`.
#### Managing Mangled Names
Just mangling names is not enough. Consider the following program:
```text {linenos=table}
defn packOne x = {
let {
data Packed a = { Pack a }
} in {
Pack x
}
}
defn packTwo x = {
let {
data Packed a = { Pack a }
} in {
Pack x
}
}
```
{{< sidenote "right" "lifting-types-note" "Lifting the data type declarations" >}}
We are actually not quite doing something like the following snippet.
The reason for this is that we don't mangle the names for types. I pointed
out this potential issue in a sidenote in the previous post. Since the size
of this post is already balooning, I will not deal with this issue here.
Even at the end of this post, our compiler will not be able to distinguish
between the two Packed types. We will hopefully get to it later.
{{< /sidenote >}} and their constructors into the global
scope gives us something like:
``` {linenos=table}
data Packed a = { Pack a }
data Packed_1 a = { Pack_1 a }
defn packOne x = { Pack x }
defn packTwo x = { Pack_1 x }
```
Notice that we had to rename one of the calls to `Pack` to be a call to
be `Pack_1`. To actually change our AST to reference `Pack_1`, we'd have
to traverse the whole tree, and make sure to keep track of definitions
that could shadow `Pack` further down. This is cumbersome. Instead, we
can mark a variable as referring to a mangled version of itself, and
access this information when needed. To do this, we add the `mangled_name`
field to the `variable_data` struct as we've seen above, and implement
the `set_mangled_name` and `get_mangled_name` methods. The former:
{{< codelines "C++" "compiler/12/type_env.cpp" 34 37 >}}
And the latter:
{{< codelines "C++" "compiler/12/type_env.cpp" 39 45 >}}
We don't allow `set_mangled_name` to affect variables that are declared
above the receiving `type_env`, and use the empty string as a 'none' value.
Now, when lifting data type constructors, we'll be able to use
`set_mangled_name` to make sure constructor calls are made correctly. We
will also be able to use this in other cases, like the translation
of local function definitions.
#### New AST Nodes
Finally, it's time for us to add new AST nodes to our language.
Specifically, these nodes are `ast_let` (for `let/in` expressions)
and `ast_lambda` for lambda functions. We declare them as follows:
{{< codelines "C++" "compiler/12/ast.hpp" 131 166 >}}
In `ast_let`, the `definitions` field corresponds to the original definitions
given by the user in the program, and the `in` field corresponds to the
expression which uses these definitions. In the process of lifting, though,
we eventually transfer each of the definitions to the global scope, replacing
their right hand sides with partial applications. After this transformation,
all the data type definitions are effectively gone, and all the function
definitions are converted into the simple form `x = f a1 ... an`. We hold
these post-transformation equations in the `translated_definitions` field,
and it's them that we compile in this node's `compile` method.
In `ast_lambda`, we allow multiple parameters (like Haskell's `\x y -> x + y`).
We store these parameters in the `params` field, and we store the lambda's
expression in the `body` field. Just like `definition_defn`,
the `ast_lambda` node maintains a separate environment in which its children
have been bound, and a list of variables that occur freely in its body. The
former is used for typechecking, while the latter is used for lifting.
Finally, the `translated` field holds the lambda function's form
after its body has been transformed into a global function. Similarly to
`ast_let`, this node will be in the form `f a1 ... an`.
The
observant reader will have noticed that we have a new method: `translate`.
This is a new method for all `ast` descendants, and will implement the
steps of moving definitions to the global scope and transforming the
program. Before we get to it, though, let's look at the other relevant
pieces of code for `ast_let` and `ast_lambda`. First, their grammar
rules in `parser.y`:
{{< codelines "text" "compiler/12/parser.y" 107 115 >}}
This is pretty similar to the rest of the grammar, so I will give this no
further explanation. Next, their `find_free` and `typecheck` code.
We can start with `ast_let`:
{{< codelines "C++" "compiler/12/ast.cpp" 275 289 >}}
As you can see, `ast_let::find_free` works in a similar manner to `ast_case::find_free`.
It finds the free variables in the `in` node as well as in each of the definitions
(taking advantage of the fact that `definition_group::find_free` populates the
given set with "far away" free variables). It then filters out any variables bound in
the `let` from the set of free variables in `in`, and returns the result.
Typechecking in `ast_let` relies on `definition_group::typecheck`, which holds
all of the required functionality for checking the new definitions.
Once the definitions are typechecked, we use their type information to
typecheck the `in` part of the expression (passing `definitions.env` to the
call to `typecheck` to make the new definitions visible).
Next, we look at `ast_lambda`:
{{< codelines "C++" "compiler/12/ast.cpp" 344 366 >}}
Again, `ast_lambda::find_free` works similarly to `definition_defn`, stripping
the variables expected by the function from the body's list of free variables.
Also like `definition_defn`, this new node remembers the free variables in
its body, which we will later use for lifting.
Typechecking in this node also proceeds similarly to `definition_defn`. We create
new type variables for each parameter and for the return value, and build up
a function type called `full_type`. We then typecheck the body using the
new environment (which now includes the variables), and return the function type we came up with.
#### Translation
Recalling the transformations we described earlier, we can observe two
major steps to what we have to do:
1. Move the body of the original definition into its own
global definition, adding all the captured variables as arguments.
2. Replace the right hand side of the `let/in` expression with an application
of the global definition to the variables it requires.
We will implement these in a new `translate` method, with the following
signature:
```C++
void ast::translate(global_scope& scope);
```
The `scope` parameter and its `add_function` and `add_constructor` methods will
be used to add definitions to the global scope. Each AST node will also
use this method to implement the second step. Currently, only
`ast_let` and `ast_lambda` will need to modify themselves - all other
nodes will simply recursively call this method on their children. Let's jump
straight into implementing this method for `ast_let`:
{{< codelines "C++" "compiler/12/ast.cpp" 291 316 >}}
Since data type definitions don't really depend on anything else, we process
them first. This amounts to simply calling the `definition_data::into_globals`
method, which itself simply calls `global_scope::add_constructor`:
{{< codelines "C++" "compiler/12/definition.cpp" 86 92 >}}
Note how `into_globals` updates the mangled name of its constructor
via `set_mangled_name`. This will help us decide which global
function to call during code generation. More on that later.
Starting with line 295, we start processing the function definitions
in the `let/in` expression. We remember how many arguments were
explicitly added to the function definition, and then call the
definition's `into_global` method. This method is implemented
as follows:
{{< codelines "C++" "compiler/12/definition.cpp" 40 49 >}}
First, this method collects all the non-global free variables in
its body, which will need to be passed to the global definition
as arguments. It then combines this list with the arguments
the user explicitly added to it, recursively translates
its body, and creates a new global definition using `add_function`.
We return to `ast_let::translate` at line 299. Here,
we determine how many variables ended up being captured, by
subtracting the number of explicit parameters from the total
number of parameters the new global definition has. This number,
combined with the fact that we added all the 'implict' arguments
to the function to the beginning of the list, will let us
iterate over all implict arguments, creating a chain of partial
function applications.
But how do we build the application? We could use the mangled name
of the function, but this seems inelegant, especially since we
alreaady keep track of mangling information in `type_env`. Instead,
we create a new, local environment, in which we place an updated
binding for the function, marking it global, and setting
its mangled name to the one generated by `global_sope`. This work is done
on lines 301-303. We create a reference to the global function
using the new environment on lines 305 and 306, and apply it to
all the implict arguments on lines 307-313. Finally, we
add the new 'basic' equation into `translated_definitions`.
Let's take a look at translating `ast_lambda` next:
{{< codelines "C++" "compiler/12/ast.cpp" 368 392 >}}
Once again, on lines 369-375 we find all the arguments to the
global definition. On lines 377-382 we create a new global
function and a mangled environment, and start creating the
chain of function applications. On lines 384-390, we actually
create the arguments and apply the function to them. Finally,
on line 391, we store this new chain of applications in the
`translated` field.
#### Compilation
There's still another piece of the puzzle missing, and
that's how we're going to compile `let/in` expressions into
G-machine instructions. We have allowed these expressions
to be recursive, and maybe even mutually recursive. This
worked fine with global definitions; instead of specifying
where on the stack we can find the reference to a global
function, we just created a new global node, and called
it good. Things are different now, though, because the definitions
we're referencing aren't _just_ global functions; they are partial
applications of a global function. And to reference themselves,
or their neighbors, they have to have a handle on their own nodes. We do this
using an instruction that we foreshadowed in part 5, but didn't use
until just now: __Alloc__.
__Alloc__ creates placeholder nodes on the stack. These nodes
are indirections, the same kind that we use for lazy evaluation
and sharing elsewhere. We create an indirection node for every
definition that we then build; when an expression needs access
to a definition, we give it the indirection node. After
building the partial application graph for an expression,
we use __Update__, making the corresponding indirection
point to this new graph. This way, the 'handle' to a
definition is always accessible, and once the definition's expression
is built, the handle correctly points to it. Here's the implementation:
{{< codelines "C++" "compiler/12/ast.cpp" 319 332 >}}
First, we create the __Alloc__ instruction. Then, we update
our environment to map each definition name to a location
within the newly allocated batch of nodes. Since we iterate
the definitions in order, 'pushing' them into our environment,
we end up with the convention of having the later definitions
closer to the top of the G-machine stack. Thus, when we
iterate the definitions again, this time to compile their
bodies, we have to do so starting with the highest offset,
and working our way down to __Update__-ing the top of the stack.
Once the definitions have been compiled, we proceed to compiling
the `in` part of the expression as normal, using our updated
environment. Finally, we use __Slide__ to get rid of the definition
graphs, cleaning up the stack.
Compiling the `ast_lambda` is far more straightforward. We just
compile the resulting partial application as we normally would have:
{{< codelines "C++" "compiler/12/ast.cpp" 394 396 >}}
One more thing. Let's adopt the convention of storing __mangled__
names into the compilation environment. This way, rather than looking up
mangled names only for global functions, which would be a 'gotcha'
for anyone working on the compiler, we will always use the mangled
names during compilation. To make this change, we make sure that
`ast_case` also uses `mangled_name`:
{{< codelines "C++" "compiler/12/ast.cpp" 242 242 >}}
We also update the logic for `ast_lid::compile` to use the mangled
name information:
{{< codelines "C++" "compiler/12/ast.cpp" 52 58 >}}
#### Fixing Type Generalization
This is a rather serious bug that made its way into the codebase
since part 10. Recall that we can only generalize type variables
that are free in the environment. Thus far, we haven't done that,
and we really should: I ran into incorrectly inferred types
in my first test of the `let/in` language feature.
We need to make our code capable of finding free variables in the
type environment. This requires the `type_mgr`, which associates
with type variables the real types they represent, if any. We
thus create methods with signatures as follows:
```C++
void type_env::find_free(const type_mgr& mgr, std::set& into) const;
void type_env::find_free_except(const type_mgr& mgr, const std::string& avoid,
std::set& into) const;
```
Why `find_free_except`? When generalizing a variable whose type was already
stored in the environment, all the type variables we could generalize would
not be 'free'. If they only occur in the type we're generalizing, though,
we shouldn't let that stop us! More generally, if we see type variables that
are only found in the same mutually recursive group as the binding we're
generalizing, we are free to generalize them too. Thus, we pass in
a reference to a `group`, and check if a variable is a member of that group
before searching it for free type variables. The implementations of the two
methods are straightforward:
{{< codelines "C++" "compiler/12/type_env.cpp" 4 18 >}}
Note that `find_free_except` calls `find_free` in its recursive call. This
is not a bug: we _do_ want to include free type variables from bindings
that have the same name as the variable we're generalizing, but aren't found
in the same scope. As far as we're concerned, they're different variables!
The two methods use another `find_free` method which we add to `type_mgr`:
{{< codelines "C++" "compiler/12/type.cpp" 206 219 >}}
This one is a bit of a hack. Typically, while running `find_free`, a
`type_mgr` will resolve any type variables. However, variables from the
`forall` quantifier of a type scheme should not be resolved, since they
are explicitly generic. To prevent the type manager from erroneously resolving
such type variables, we create a new type manager that does not have
these variables bound to anything, and thus marks them as free. We then
filter these variables out of the final list of free variables.
Finally, `generalize` makes sure not to use variables that it finds free:
{{< codelines "C++" "compiler/12/type_env.cpp" 68 81 >}}
#### Putting It All Together
All that's left is to tie the parts we've created into one coherent whole
in `main.cpp`. First of all, since we moved all of the LLVM-related
code into `global_scope`, we can safely replace that functionality
in `main.cpp` with a method call:
{{< codelines "C++" "compiler/12/main.cpp" 121 132 >}}
On the other hand, we need top-level logic to handle `definition_group`s.
This is pretty straightforward, and the main trick is to remember to
update the function's mangled name. Right now, depending on the choice
of manging algorithm, it's possible even for top-level functions to
have their names changed, and we must account for that. The whole code is:
{{< codelines "C++" "compiler/12/main.cpp" 52 62 >}}
Finally, we call `global_scope`'s methods in `main()`:
{{< codelines "C++" "compiler/12/main.cpp" 148 151 >}}
That's it! Please note that I've mentioned or hinted at minor changes to the
codebase. Detailing every single change this late into the project is
needlessly time consuming and verbose; Gitea reports that I've made 677
insertions into and 215 deletions from the code. As always, I provide
the [source code for the compiler](https://dev.danilafe.com/Web-Projects/blog-static/src/branch/master/code/compiler/12), and you can also take a look at the
[Gitea-generated diff](https://dev.danilafe.com/Web-Projects/blog-static/compare/1905601aaa96d11c771eae9c56bb9fc105050cda...21851e3a9c552383ee8c4bc878ea06e7d28c333e)
at the time of writing. If you want to follow along, feel free to check
them out!
### Running Our Programs
It's important to test all the language features that we just added. This
includes recursive definitions, nested function dependency cycles, and
uses of lambda functions. Some of the following examples will be rather
silly, but they should do a good job of checking that everything works
as we expect. Let's start with a simple use of a recursive definition
inside a `let/in`. A classic definition in that form is of `fix`
(the fixpoint combinator):
```Haskell
fix f = let x = f x in x
```
This defines `x` to be `f x`, which by substitution becomes `f (f x)`, and then
`f (f (f x))` and so on. The fixpoint combinator allows one to write a
recursive function that doesn't use its own name in the body. Rather,
we write a function expecting to receive 'itself' as a value:
```Haskell
fix :: (a -> a) -> a
factRec :: (Int -> Int) -> Int -> Int
factRec f x = if x == 0 then 1 else x * f x
fact :: Int -> Int
fact = fix factRec
```
Notice that `factRec` doesn't reference itself, but rather takes
as argument a function it expects to be 'factorial' called `f`,
and uses that in its recursive case. We can write something similar
in our language, perhaps to create an infinite list of ones:
{{< codeblock "text" "compiler/12/examples/fixpoint.txt" >}}
We want `sumTwo` to take the first two elements from the list,
and return their sum. For an infinite list of ones, we expect
this sum to be equal to 2, and it is:
```
Result: 2
```
Next, let's try to define a function which has a mutually recursive pair
of definitions inside of a `let/in`. Let's also make these expressions
reference a function from the global scope, so that we know our
dependency tracking works as expected:
{{< codeblock "text" "compiler/12/examples/letin.txt" >}}
Here, we have a function `mergeUntil` which, given two lists
and a predicate, combines the two lists as long as
the predicate returns `True`. It does so using a convoluted
pair of mutually recursive functions, one of which
unpacks the left list, and the other the right. Each of the
functions calls the global function `if`. We also use two
definitions inside of `main` to create the two lists we're
going to merge. The compiler outputs the following (correct)
types:
```
const: forall bb bc . bc -> bb -> bc
if: Bool* -> List* Int* -> List* Int* -> List* Int*
main: Int*
mergeUntil: List* Int* -> List* Int* -> (Int* -> Bool*) -> List* Int*
sum: List* Int* -> Int*
```
And the result is 21, as would be expected from the sum of the numbers 1-6:
```
Result: 21
```
Let's try lambda functions now. We can try use them for a higher-order function
like `map`:
{{< codeblock "text" "compiler/12/examples/lambda.txt" >}}
In this example, we first double every element in the list, then square it,
and finally take the sum. This should give us 4+16+36 = 56, and so it does:
```
Result: 56
```
Finally, let's do some magic with a locally-declared data type. We'll make a
"packer" that creates a wrapped instance of a type, `Packed a`. Since the
constructor of this data type is not globally visible, it's not possible
to get the value back out, except by using an 'unpacking' function that
we provide:
{{< codeblock "text" "compiler/12/examples/packed.txt" >}}
Here, the `packer` definition returns a pair of the 'packing'
and 'unpacking' functions. The 'packing' function simply applies
the consntructor of `Packed` to its argument, while the 'unpacking'
function performs pattern matching (which is possible since the
data type is still in scope there). We expect `unpack (pack 3)` to
return 3, and it does:
```
Result: 3
```
Trying to pattern match, though, doesn't work, just like we would want!
This is enough to convince me that our changes do, indeed, work! Of
the 'major' components that I wanted to cover, only __Input/Output__
remains! Additionally, a [lobste.rs](https://lobste.rs) user suggested
that we also cover namespacing, and perhaps we will look into that as well.
Before either of those things, though, I think that I want to go through
the compiler and perform another round of improvements, similarly to
[part 4]({{< relref "04_compiler_improvements" >}}). It's hard to do a lot
of refactoring while covering new content, since major changes need to
be explained and presented for the post to make sense.
I do this in [part 13]({{< relref "13_compiler_cleanup/index.md" >}}) - cleanup.
I hope to see you there!