7.6

## First-class Functions

Based on these notes by Ben Lerner. Various small changes and corrections.

### 1First-class Functions

In Defining functions, we introduced the ability for our programs to define functions that we could then call in other expressions in our program. Our programs were a sequence of function definitions, followed by one main expression. This notion of a program was far more flexible than we had before, and lets us define many computations we simply could not previously do. But it is distinctly unsatisfying: functions are second-class entities in our language, and can’t be used the same way as other values in our programs.

We know from other courses, and indeed even from writing compilers in ML, that higher-order functions — functions whose arguments can be functions — are very useful notions to have. Let’s consider the most trivial higher-order program:

(def (applyToFive it)
(it 5))

(def (incr x)
(+ x 1))

(applyToFive incr)

Do Now!

What errors currently get reported for this program?

Because it is a parameter to the first function, our compiler will complain that it is not defined as a function, when used as such on line 2. Additionally, because incr is defined as a function, our compiler will complain that it can’t be used as a parameter on the last line. We’d like to be able to support this program, though, and others more sophisticated. Doing so will bring in a number of challenges, whose solutions are detailed and all affect each other. Let’s build up to those programs, incrementally.

### 2Reminder: How are functions currently compiled?

Let’s simplify away the higher-order parts of the program above, and look just at a basic function definition. The following program:

(def (incr x)
(+ x 1))

(incr 5)

is compiled to:

incr:
push RBP            ;; stack frame management
mov RBP, RSP

mov RAX, [RBP + 16] ;; get param

mov RSP, RBP        ;; undo stack frame
pop RBP
ret                 ;; exit

our_code_starts_here:
push RBP            ;; stack frame management
mov RBP, RSP

push 10             ;; push (encoded) 5
call incr           ;; call function
add RSP, 8          ;; remove arguments

mov RSP, RBP        ;; undo stack frame
pop RBP
ret                 ;; exit

This compilation is a pretty straightforward translation of the code we have. What can we do to start supporting higher-order functions?

### 3The value of a function — Attempt #1

#### 3.1Passing in functions

Going back to the original motivating example, the first problem we encounter is seen in the first and last lines of code.

(def (applyToFive it)
(it 5))

(def (incr x)
(+ x 1))

(applyToFive incr)

Functions receive values as their parameters, and function calls push values onto the stack. So in order to “pass a function in” to another function, we need to answer the question, what is the value of a function? In the assembly above, what could possibly be a candidate for the value of the incr function?

A function, as a standalone entity, seems to just be the code that comprises its compiled body. We can’t conveniently talk about the entire chunk of code, though, but we don’t actually need to. We really only need to know the “entrance” to the function: if we can jump there, then the rest of the function will execute in order, automatically. So one prime candidate for “the value of a function” is the address of its first instruction. Annoyingly, we don’t know that address explicitly, but fortunately, the assembler helps us here: we can just use the initial label of the function, whose name we certainly do know.

In other words, we can compile the main expression of our program as:

our_code_starts_here:
push RBP           ;; stack frame management
mov RBP, RSP

push incr          ;; push the start label of incr
call applyToFive   ;; call function
add RSP, 8         ;; remove arguments

mov RSP, RBP       ;; undo stack frame
pop RBP
ret                ;; exit

This might seem quite bizarre: how can we push a label onto the stack? Doesn’t push require that we push a value — either a constant, or a register’s value, or some word of memory? In fact it is no more and no less bizarre than calling a label in the first place: the assembler replaces those named labels with the actual addresses within the program, and so at runtime, they’re simply normal word values representing memory addresses.

#### 3.2Using function arguments

Do Now!

The compiled code for applyToFive looks like this:

applyToFive:
push RBP            ;; stack frame management
mov RBP, RSP

mov RAX, [RBP + 16] ;; get the param
push ????           ;; push the argument to it
call ????           ;; call it
add RSP, 8          ;; remove arguments

mov RSP, RBP        ;; undo stack frame
pop RBP
ret                 ;; exit

Fill in the questions to complete the compilation of applyToFive.

The parameter for it is simply 5, so we push 10 onto the stack, just as before. The function to be called, however, isn’t identified by its label: we already have its address, since it was passed in as the argument to applyToFive. Accordingly, we call RAX in order to find and call our function. Again, this generalizes the syntax of call instructions slightly just as push was generalized: we can call an address given by a register, instead of just a constant.

#### 3.3Victory!

We can now pass functions to functions! Everything works exactly as intended.

Do Now!

Tweak the example program slightly, and cause it to break. What haven’t we covered yet?

### 4The measure of a function — Attempt #2

Just because we use a parameter as a function doesn’t mean we actually passed a function in as an argument. If we change our program to applyToFive(true), our program will attempt to apply true as a function, meaning it will try to call 0xFFFFFFFFFFFFFFFF, which isn’t likely to be a valid address of a function.

As a second, related problem: suppose we get bored of merely incrementing values by one, and generalize our program slightly:

(def (applyToFive it)
(it 5))

(+ x y))

(applyToFive add)

Do Now!

What happens now?

Let’s examine the stack very carefully. When our program starts, it pushes add onto the stack, then calls applyToFive. (The colors indicate which functions control the data on the stack, while the brackets along the side indicate which function uses the data on the stack; explaining why they don’t quite align at function-argument positions.) That function in turn pushes 10 onto the stack, and calls it (i.e. the address currently stored in RAX): But look at the bracketing for add! It needs two arguments, but receives only one. So it adds 5 (encoded as 10) to the saved RBP, since as far as it knows that stack location is where its second parameter should be.

We had eliminated both of these problems before via well-formedness checking: our function-definition environment knew about every function and its arity, and we could check every function application to ensure that a well-known function was called, with the correct number of arguments were passed. But now that we can pass functions around dynamically, we can’t know statically whether the arities are correct, and can’t even know whether we have a function at all!

We don’t know anything about precisely where a function’s code begins, so there’s no specific property we could check about the value passed in to determine if it actually is a function. But in any case, that value is insufficient to encode both the function and its arity. Fortunately, we now have a technique for storing multiple pieces of data as a single value: tuples. So our second candidate for “the value of a function” is a tuple containing the function’s arity and start address. This isn’t quite right either, since we wouldn’t then be able to distinguish actual tuples from “tuples-that-are-functions”.

So we choose a new tag value, say 0x5, distinct from the ones used so far, to mark these function values. Even better: we now have free rein to separate and optimize the representation for functions, rather than hew completely to the tuple layout. As one immediate consequence: we don’t need to store the tuple length — it’s always 2, namely the function arity and the function pointer. So we might as well instead store the function arity in the header word where the tuple length used to be: since we’ll have a tag to tell us “this is a function value, not a tuple”, we won’t misinterpret the header word by mistake.

Do Now!

Revise the compiled code of applyToFive to assume it gets one of the new tuple-like values.

The pseudocode for calling a higher-order function like this is roughly:

mov RAX, <the function tuple>  ;; load the intended function
<check-tag RAX, 0x5>           ;; ensure it has the right tag
sub RAX, 5                     ;; untag the value
<check-arity [RAX], num-args>  ;; the word at [RAX] stores the arity
<push all the args>            ;; set up the stack
call [RAX+8]                   ;; the second word stores the function address
add RSP, <8 * num-args>        ;; finish the call

Now we just need to create these tuples.

Exercise

Revise the compiled code above to allocate and tag a function value using this new scheme, instead of a bare function pointer.

### 5A function by any other name — Attempt #3

We now have a disparity between “normal” function calls, where we know the name (and arity) comes from a top-level declaration in the source program, and “higher-order” function calls, where the function to be called comes in as a parameter, and is represented as a tuple.

Because of this distinction, we don’t have any good, uniform way to handle compiling function calls. In particular, we don’t have an obvious place in our compilation to create those tuples.

What if we revise our language, to make functions be just another expression form, rather than a special top-level form? We’ve seen these in other languages: we call them lambda expressions, and they appear in pretty much all major languages:

 Language Lambda syntax Haskell \(x1,...,xn) -> e Ocaml fun (x1,...,xn) -> e Javascript (x1,...,xn) => { return e; } C++ [&](x1,...,xn){ return e; }

We can rewrite our initial example as

(let (applyToFive (lambda (it) (it 5))) in
(let (incr (lambda (x) (+ x 1))) in
(incr applyToFive)))

Now, all our functions are defined in the same manner as any other let-binding: they’re just another expression, and we can simply produce the function values right then, storing them in let-bound variables as normal.

Aside: Another valid option is to keep first-class and second-class functions separate, and use a different syntax for application in order to differentiate which function call semantics we are expecting. An example of this approach is Common Lisp, where there are two namespaces, one for functions and one for variables. In Common Lisp, to apply a lambda bound to the name foo in the variables namespace, one needs to write (funcall foo 1 2). Dually, to pass a defined function bar as argument, one needs to write (function bar), as in (map (function bar) my_list). There are many possible designs here, pick the one you want (and document it).

Let’s come back to lambdas and try compiling a simplified version of the code above:

(let (incr (lambda (x) (+ x 1)))
(incr 5))

Our compiled output will look something like this:

our_code_starts_here:
push RBP
mov RBP, RSP

incr:
push RBP
mov RSP, RBP

mov RAX, [RBP+8]

mov RSP, RBP
pop RBP
ret

mov RAX, R15   ;; allocate a function tuple
or RAX, 0x5    ;; tag it as a function tuple
mov [R15+0], 1 ;; set the arity of the function
mov [R15+8], incr

mov [RBP-8], RAX ;; (let (incr ...) ...

mov RAX, [RBP-8]
<check that RAX is tagged 0x5>
sub RAX, 5
<check that RAX expects 1 argument>
push 10
call [RAX+8]

move RSP, RBP
pop RBP
ret

Do Now!

What’s wrong with this code?

Our program will start executing at our_code_starts_here, and flows straight into the code for incr, even though it hasn’t been called! We seem to have left out a crucial part of the semantics of functions: while a function is defined by its code, that code should not run until it’s called: lambda-expressions are inert values.

Do Now!

What simple code-generation tweak can we use to fix this?

On the one hand, the code shouldn’t be run. On the other, we have to emit the code somewhere. There are two possible solutions here:

• We can transform our program even further, to somehow lift all the lambdas out from the innards of other functions so that we regain the “every function lives at the top level” structure of our old code. This approach, called lambda-lifting, works well, but is overkill for our purposes for now.

• Another is simply to label the end of the function, and just add a jmp end_label instruction before the initial label. We then bypass the code of the function when we’re “defining” it, but when we call it, we skip the jmp and start right at the first instruction of the code.

Exercise

Compile the original example to assembly by hand.

### 6“Objects in mirror may be closer than they appear” — Attempt #4

Our running example annoyingly hard-codes the increment operation. Let’s generalize:

(let (add (lambda (x) (lambda (y) (+ x y))))
(let (applyToFive (lambda (it) (it 5)))
(tup (applyToFive incr) (applyToFive add5))))))

Exercise

What does this program produce? What goes wrong here? Draw the stack demonstrating the problem.

Our representation of functions cannot distinguish incr from add5: they have the same arity, and point to the same function. But they’re clearly not the same function! This is a problem of scope. How can we distinguish these two functions?

#### 6.1Bound and free variables

What does incr actually evaluate to? A function-tuple (1, <code>) where the code is the compiled form of (lambda (y) (+ x y)). How exactly does that expression get compiled? When we compile the expression (+ x y), we have an environment where x is mapped to “the first function parameter”, and so is y. In other words, this expression gets compiled to the same thing as (+ y y) which is certainly not the right thing! This is similar to the problem we had a while ago, where we needed to allocate distinct stack slots for all local variables, but it is more insidious here: x and y really are the first function parameters of their respective functions, but within the inner lambda, those descriptions come into conflict.

Define a variable x as bound within an expression e if

• x appears on the left side of a let-binding, or

• e is a lambda expression and x appears as one of its parameters

Define a variable to be free if it is not bound. For instance, in

(let (x (lambda (m) (let (t m) (+ x t))))
(+ x y))

• y is certainly free within the entire expression: there are no bindings for it at all.

• x is free within the lambda: its only binding appears outside that lambda

• The uses of m and t are bound, by the lambda’s parameter and by the inner let-binding, respectively.

Now we can see the problem with our add5 and incr example: x is free within the lambdas for those two functions, but our compiled code does not take that into account. We can generalize this problem easily enough: our compilation of all free variables is broken.

#### 6.2Computing the set of free variables

We need to know exactly which variables are free within an expression, if we want to compile them properly. This can be subtle to get right: because of shadowing, not every identifier that’s spelled the same way is in fact the same name. (We saw this a few lectures ago when we discussed alpha-equivalence and the safe renaming of variables.) It’s easy to define code that appears right, but it’s tricky to convince ourselves that the code in fact is correct.

Do Now!

Define a function freeVars that computes the set of free variables of a given expression.

Now what?

#### 6.3Using free variables properly: achieving closure

We know from using lambdas in other languages what behavior we expect from them: their free variables ought to take on the values they had at the moment the lambda was evaluated, rather than the moment the lambda’s code was called.1Think carefully about what your intuition is here, when a free variable is mutable... We say that we want lambdas to close over their free variables, and we describe the value of a function as a closure (rather than the awkward “function-tuple” terminology we’ve had so far). To accomplish this, we clearly need to store the values of the free variables in a reliable location, so that the compiled function body can find them when needed...and so that distinct closures with the same code but different closed-over values can behave distinctly! The natural place to store these values is in our tuple, after the function-pointer. We might consider also storing the number of closed-over variables; we’ll store that between the function-pointer and the closed-over values.

Do Now!

What language feature (that we may or may not have yet) might want to know how many closed-over variables are in our closure?

This leads to our latest (and final?) representation choice for compiling first-class functions.

Let’s work through a short example:

(let (five 5)
(let (applyToFive (lambda (it) (it five)))
(let (incr (lambda (x) (+ x 1)))
(applyToFive incr))))

First, let’s focus on the compilation of the let-binding of applyToFive: our closure should be a 4-tuple (arity = 1, code = applyToFive, size = 1, five = 5), where I’ve labelled the components for clarity.

  ...
mov RAX, 10
mov [RBP-8], RAX ;; (let (five 5) ...

jmp applyToFive_end
applyToFive:
...
applyToFive_end:

mov [R15+0], 1           ;; set the arity of the function
mov [R15+8], applyToFive ;; set the code pointer
mov [R15+16], 1          ;; number of closed-over variables
mov RAX, [RBP-8]         ;; load five
mov [R15+24], RAX        ;; store it in the closure
mov RAX, R15             ;; start allocating a closure
add RAX, 0x5             ;; tag it as a closure
add R15, 32              ;; bump the heap pointer, maintaining 16-byte alignment

mov [RBP-8], RAX         ;; (let (applyToFive ...) ...
...

Do Now!

This example shows only a single closed-over variable. What ambiguity have we not addressed yet, for closing over multiple variables?

### 7Implementing the new compilation

#### 7.1Scope-checking

Exercise

What should the new forms of well-formedness and scope checking actually do? What needs to change?

#### 7.2Compiling function bodies

Now we just need to update the compilation of the function body itself, to look for closed-over variables in the correct places. We can codify our access strategy by updating the environment we use when we compile a function body. We have two options here:

1. We can repeatedly access each free variable from the appropriate slot of the closure

2. We can unpack the closure as part of the function preamble, copying the values onto the stack as if they were let-bound variables, and offsetting our compilation of any local let-bound variables by enough slots to make room for these copies.

Exercise

What are some of the design tradeoffs of these two approaches?

Note that, either way, the compiled function body needs access at runtime to the closure itself in order to retrieve values from out of it!

#### 7.3Compiling function calls

We need to change how we compile function applications, too, in order to make our compilation of closures work. Let’s agree to change our calling signature, such that the first argument to every function call is the closure itself. In other words, the compilation of function calls will now look like:

1. Retrieve the function value, and check that it’s tagged as a closure.

2. Check that the arity matches the number of arguments being applied.

3. Push all the arguments.

4. Push the closure itself.

5. Call the code-label in the closure.

6. Pop the arguments and the closure.

#### 7.4Revisiting compiling function bodies

Our function bodies will now be compiled as something like

1. Compute the free-variables of the function.

2. Update the environment:

• All the arguments are now offset by one slot from our earlier compilation (since the first argument will be the closure itself).

• Depending on the strategy we choose: all the free variables are mapped to the first few local-variable slots, or to specific offsets with respect to the closure.

3. Compile the body in the new environment

4. If we pick the local-variable strategy, we must produce compiled code that, after the stack management and before the body, reads the saved free-variables out of the closure (which is passed in as the first function parameter), and stores them in the reserved local variable slots.

5. The closure itself is a heap-allocated tuple (arity, code-pointer, N, free-var1, ... free-varN).

#### 7.5Complete worked example

Let’s try compiling the following program (which could be the result of an ANF transformation and desugaring of def into let-lambda):

(let (foo (lambda w x y z)
(lambda (a)
(let (temp1 (+ a x))
(+ temp1 z))))
(let (temp2 (foo 1 2 3 4))
(temp2 5)))

Looking at the outermost program, we see two bindings (for foo and temp2), so we allocate stack slots for them, leading to an environment for code-generation of

foo   ==> [RBP - 8]
temp2 ==> [RBP - 16]

Within the lambda for foo, we initially see a code-generation environment containing just the four arguments. But remember that we changed our calling convention, and there is now an implicit “self” argument as the very first argument:

self ==> [RBP + 16]
w    ==> [RBP + 24]
x    ==> [RBP + 32]
y    ==> [RBP + 40]
z    ==> [RBP + 48]

(Of course, the schema above must be updated if we’re using the x86-64 calling convention, since then all five arguments will come from registers.)

Within the innermost lambda, our code-generation environment is trickier. For scope checking, we have the five names w, x, y, z and a in scope. Of those, the only free variables in the lambda are x and z. So our code-generation environment needs to include them, as well as the arguments of the lambda. As a first draft, our environment will be

self ==> [RBP + 16]
a    ==> [RBP + 24]
x    ==> [self + 24]
z    ==> [self + 32]

Accessing x or z requires two memory lookups, first to load the closure and second to offset into it; this double-indirection is one reason why one might choose to unpack the closure within the function body. In that case, the environment for the lambda would to actually be

self ==> [RBP + 16]
a    ==> [RBP + 24]
x    ==> [RBP - 8]   ; free variables are
z    ==> [RBP - 16]  ; on the stack

We’ll need to compile the lambda’s body with this environment. We then must be careful to account for this initial use of the stack, so that we bind temp1 to [RBP - 24]. The compiled code of our innermost lambda will be:

    inner_lambda:
;; prologue
push RBP
mov RBP, RSP
;; unpack the closure
sub RSP, 16         ;; reserve space on the stack for closed-over vars
mov R11, [RBP + 16] ;; \ load the self argument
sub R11, 0x5        ;; / and untag it
mov RAX, [R11 + 24] ;; \ load x from closure
mov [RBP - 8], RAX  ;; / into its correct stack slot
mov RAX, [R11 + 16] ;; \ load z from closure
mov [RBP - 16], RAX ;; / into its correct stack slot
;; actual function body
sub RSP, 8          ;; reserve space on the stack for locals
mov RAX, [RBP + 24] ;; \
add RAX, [RBP - 8]  ;; | (let (temp1 (+ a x)) ...)
mov [RBP - 24], RAX ;; /
mov RAX, [RBP - 24] ;; \ (+ temp1 z)
add RAX, [RBP - 16] ;; /
;; epilogue
mov RSP, RBP
pop RBP
ret
inner_lambda_end:

Note that this is not a closure yet! It’s merely the code of our lambda. To generate a closure, we need to surround it by additional work. This work takes place in the environment for foo:

    ;; skip over the code for the lambda
jmp inner_lambda_end
inner_lambda:
... everything above ...
inner_lambda_end:
;; start filling in the closure information
mov [R15 + 0 ], 1            ;; arity
mov [R15 + 8 ], inner_lambda ;; code pointer
mov [R15 + 16], 2            ;; # of free variables
mov RAX, [RBP + 32]          ;; \ copy x from argument
mov [R15 + 24], RAX          ;; / into closure
mov RAX, [RBP + 40]          ;; \ copy z from argument
mov [R15 + 32], RAX          ;; / into closure
;; start creating the closure value
mov RAX, R15                 ;; \ create the closure
add R15, 48                  ;; update heap pointer, keeping 16-byte alignment
;; now RAX contains a proper closure value

All this code lives inside the compiled body of foo:

  foo:
;; prologue
push RBP
mov RBP, RSP
;; unpack the closure -- nothing to do here
;; actual function body
;; skip over the code for the lambda
... everything above ...
;; now RAX contains a proper closure value
;; since the inner lambda is in tail position,
;; RAX is our return value, so we're done.
;; epilogue
mov RSP, RBP
pop RBP
ret
foo_end:

Once again, note that this is not a closure yet! It’s merely the code of our foo lambda. To generate a closure, we need to surround it by additional work. This work takes place in the environment for our_code_starts_here:

  ;; skip over the code for foo
jmp foo_end
foo:
... everything above ...
foo_end:
;; start filling in the closure information
mov [R15 + 0 ], 4            ;; arity
mov [R15 + 8 ], foo          ;; code pointer
mov [R15 + 16], 0            ;; # of free variables
;; start creating the closure value
mov RAX, R15                 ;; \ create the closure
add R15, 32                  ;; update heap pointer, keeping 16-byte alignment
;; now RAX contains a proper closure value

All this code lives inside the compiled body of our_code_starts_here:

our_code_starts_here:
;; prologue
push RBP
mov RBP, RSP
;; skip over the code for foo
... everything above ...
;; now RAX contains a proper closure value
mov [RBP - 8], RAX          ;; (let (foo (lambda ...)) ...)
;; start calling (foo 1 2 3 4)
mov RAX, [RBP - 8]
mov R11, RAX              ;; load up the value
and R11, 0x7              ;; mask off the non-tag bits
cmp R11, 0x5              ;; \ check for desired tag
jne error_not_closure     ;; /
sub RAX, 0x5              ;; untag
cmp [RAX + 0, 8]          ;; \ check arity
jne error_wrong_arity     ;; /
push 0x8                  ;; \
push 0x6                  ;; | push arguments
push 0x4                  ;; |
push 0x2                  ;; /
push [RBP - 8]            ;; push closure itself
call [RAX + 8]            ;; make the call
add RSP, 40               ;; pop arguments
mov [RBP - 16], RAX       ;; (let (temp2 (foo 1 2 3 4)) ...)
;; start calling (temp2 5)
mov R11, RAX              ;; load up the value
and R11, 0x7              ;; mask off the non-tag bits
cmp R11, 0x5              ;; \ check for desired tag
jne error_not_closure     ;; /
sub RAX, 0x5              ;; untag
cmp [RAX + 0, 8]          ;; \ check arity
jne error_wrong_arity     ;; /
push 0x10                 ;; push argument
push [RBP - 16]           ;; push closure itself
call [RAX + 8]            ;; make the call
pop 16                     ;; pop arguments
;; epilogue
mov RSP, RBP
pop RBP
ret
our_code_starts_here_end:

Note that we have to be a bit careful about our_code_starts_here: if it gets compiled as if it were a closure, then when we call it from main we have to be careful to pass itself in as the first argument: our_code_starts_here(our_code_starts_here, ... any further arguments). Or, we could special-case it slightly, and not treat it as following the standard C calling convention rather than our revised closure-tolerant calling convention.

#### 7.6Revisiting our runtime

If we unify second-class functions (defined with def at the top-level) and first-class functions (defined with lambda), we will soon realize that we have a problem: all our runtime-provided functions, like print, now no longer work with our new calling signature: we provide one too many arguments. Worse, before we could identify all functions by their starting label, and runtime-provided functions had starting labels too. Now, we need closures, but our runtime-provided functions don’t have such things.

Now that our calling convention for our own functions has changed, perhaps it makes sense to distinguish between calls to our own functions, versus calls to runtime-provided “native” functions. For that we need to recognize when to produce a “native call”, and when to produce a “closure-friendly call”.

The alternative design, which keeps both second-class functions as we had them (and encompass both def functions and external functions) and lambdas separate also needs a way to differentiate the two kinds of calls (recall the note about Common Lisp above).

Whatever option you pick, some differentiation must be made. Again, choose the design you prefer, and document it!

Exercise

Think about implementing different approaches. Which do you find (or, anticipate to be) easier, and why? What are the pros and cons? For instance, which of them actually allow passing native functions as arguments to higher-order functions? Why or why not?

We can now pass functions with free-variables to functions! Everything works exactly as intended.

### 8Recursion

Note that if we choose to desugar def as let-lambda, and we try even a simple recursive function — something that worked with our previous top-level function definitions — we run into a problem. Because we now only have let-bindings and anonymous lambdas, we have no way to refer to the function itself from within the function. We’ll get a scope error during well-formedness checking; such a program wouldn’t even make it to compilation.

(let (fac (lambda (n)
(if (< n 1)
1
(* n (fac (- n 1)))))) # ERROR: fac is not in scope
(fac 5))

This is precisely why ML has a let rec construct (called letrec in Scheme): we need to inform our compiler that the name being bound should be considered as bound within the binding itself.2That’s a twisty description, but then again, we are dealing with recursion... But not just every expression can be considered recursively bound:

let rec x = x + x in "this is nonsense"

Pretty much, only functions can be let-rec-bound: they are the only form of value we have that does not immediately trigger computation that potentially uses the name-being-bound. They are deferred computation, awaiting being called.3Lazy languages, like Haskell, allow some other recursive definitions that are not explicitly functions, such as let ones = 1 : ones. Essentially, all values in lazy languages are semantically equivalent to zero-argument functions, which are only called at most once (when they are truly needed), and whose values are cached, so they’re never recomputed; these implicit thunks are how these recursive definitions can be considered well-formed. The particular example above, though, of let rec x = x + x, would not work in Haskell either: addition requires its arguments to be fully evaluated, so laziness or not, x is still not yet defined by the time it is needed. We have a choice: we can introduce a new syntactic form for these let-rec-bound variables, or we can introduce a new well-formedness check.

Exercise

Extend the compilation above to work for recursive functions. Hint: the locations of free variables is going to change, to make room for something else.

Note that if we keep defined functions as second class, we have the benefit of supporting recursion for these functions. That of course does not solve the issue of defining recursive lambdas. Tradeoffs...

We can now pass functions with free-variables to functions and to themselves! Everything works exactly as intended.

1Think carefully about what your intuition is here, when a free variable is mutable...

2That’s a twisty description, but then again, we are dealing with recursion...

3Lazy languages, like Haskell, allow some other recursive definitions that are not explicitly functions, such as let ones = 1 : ones. Essentially, all values in lazy languages are semantically equivalent to zero-argument functions, which are only called at most once (when they are truly needed), and whose values are cached, so they’re never recomputed; these implicit thunks are how these recursive definitions can be considered well-formed. The particular example above, though, of let rec x = x + x, would not work in Haskell either: addition requires its arguments to be fully evaluated, so laziness or not, x is still not yet defined by the time it is needed.