blog-static/content/blog/12_compiler_let_in_lambda/index.md

---
title: Compiling a Functional Language Using C++, Part 12 - Let/In and Lambdas
date: 2020-04-20T20:15:16-07:00
tags: ["C and C++", "Functional Languages", "Compilers"]
description: "In this post, we extend our language with let/in expressions and lambda functions."
draft: true
---

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` expressions 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 examples. The most 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 <code>x</code> 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
<a href="https://en.wikipedia.org/wiki/Tacit_programming#Functional_programming">point-free</a>
programming style may be slightly twitching right now, the words <em>eta reduction</em> swirling in their mind. What do you know, <code>fold</code>-based <code>sum</code> 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 <code>safeTail</code>.
<br><br>
It's not that these things are <em>impossible</em> 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 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. <br><br>
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`!

### 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 so allow for the changes 
we're going to make; then, we can implement `let/in` expressions; finally,
we'll tackle 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 function name). This shows in
our code for `find_free`. Consider, for example, this segment:

{{< 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<std::string>& 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" 13 22 >}}

The `visibility` enum is defined as follows:

{{< codelines "C++" "compiler/12/type_env.hpp" 10 10 >}}

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" 31 44 >}}

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 <em>quite</em> 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 <code>Packed</code> 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 the `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 quickly see the parsing
rules for `ast_let` and `ast_lambda`:

{{< 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.

{{< todo >}}
Explain typechecking for lambda functions and let/in expressions.
{{< /todo >}}

{{< todo >}}
Explain free variable detection for lambda functions and let/in expressions.
{{< /todo >}}

#### Translation
While collecting all of the definitions into a global list, we can
also do some program transformations. Let's return to our earlier example:

```Haskell {linenos=table}
fourthPower x = square * square
    where
        square = x * x
```

We said it should be translated into something like:

```Haskell {linenos=table}
fourthPower x = square * square
    where square = square' x
square' x = x * x
```

In our language, the original program above would be:

```text {linenos=table}
defn fourthPower x = { 
    let {
        defn square = { x * x }
    } in {
        square * square
    }
}
```

And the translated version would be:

```text {linenos=table}
defn fourthPower x = { 
    let {
        defn square = { square' x }
    } in {
        square * square
    }
}
defn square' x = { x * x }
```

Setting aside for the moment the naming of `square'` and `square`, we observe
that we want to perform the following operations:

1. Move the body of the original definition of `square` 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.