In part 10, we managed to get our compiler to accept functions that were polymorphically typed. However, a piece of the puzzle is still missing: while our functions can handle values of different types, the same cannot be said for our data types. This means that we cannot construct data structures that can contain arbitrary types. While we can define and use a list of integers, if we want to also have a list of booleans, we must copy all of our constructors and define a new data type. Worse, not only do we have to duplicate the constructors, but also all the functions that operate on the list. As far as our compiler is concerned, a list of integers and a list of booleans are entirely different beasts, and cannot be operated on by the same code.
To make polymorphic data types possible, we must extend our language (and type system) a little. We will now allow for something like this:
data List a = { Nil, Cons a List }
In the above snippet, we are no longer declaring a single type, but a collection
of related types, parameterized by a type a
. Any type can take the place
of a
to get a list containing that type of element.
Then, List Int
is a type,
as is List Bool
and List (List Int)
. The constructors in the snippet also
get polymorphic types:
When you call Cons
, the type of the resulting list varies with the type of element
you pass in. The empty list Nil
is a valid list of any type, since, well, it’s
empty.
Let’s talk about List
itself, now. I suggest that we ponder the following table:
$\text{List}$  $\text{Cons}$ 

$\text{List}$ is not a type; it must be followed up with arguments, like $\text{List} \; \text{Int}$.  $\text{Cons}$ is not a list; it must be followed up with arguments, like $\text{Cons} \; 3 \; \text{Nil}$. 
$\text{List} \; \text{Int}$ is in its simplest form.  $\text{Cons} \; 3 \; \text{Nil}$ is in its simplest form. 
$\text{List} \; \text{Int}$ is a type.  $\text{Cons} \; 3 \; \text{Nil}$ is a value of type $\text{List} \; \text{Int}$. 
I hope that the similarities are quite striking. I claim that
List
is quite similar to a constructor Cons
, except that it occurs
in a different context: whereas Cons
is a way to create values,
List
is a way to create types. Indeed, while we call Cons
a constructor,
it’s typical to call List
a type constructor.
We know that Cons
is a function which
assigns to values (like 3
and Nil
) other values (like Cons 3 Nil
, or [3]
for
short). In a similar manner, List
can be thought of as a function
that assigns to types (like Int
) other types (like List Int
). We can
even claim that it has a type:
[note:
When your type constructors take as input not only other types but also values
such as 3
, you've ventured into the territory of
dependent types.
This is a significant step up in complexity from what we'll be doing in this
series. If you're interested, check out
Idris (if you want to know about dependent types
for functional programming) or Coq (to see how
propositions and proofs can be encoded in a dependently typed language).
]
our type constructors will only take zero or more types as input, and produce
a type as output. In this case, writing $\text{Type}$ becomes quite repetitive,
and we will adopt the convention of writing $*$ instead. The types of such
constructors are called kinds.
Let’s look at a few examples, just to make sure we’re on the same page:
 The kind of $\text{Bool}$ is $*$: it does not accept any type arguments, and is a type in its own right.
 The kind of $\text{List}$ is $*\rightarrow *$: it takes one argument (the type of the things inside the list), and creates a type from it.
 If we define a pair as
data Pair a b = { MkPair a b }
, then its kind is $* \rightarrow * \rightarrow *$, because it requires two parameters.
As one final observation, we note that effectively, all we’re doing is tracking the arity of the constructor type.
Let’s now enumerate all the possible forms that (mono)types can take in our system:
 A type can be a placeholder, like $a$, $b$, etc.
 A type can be a type constructor, applied to [note: It is convenient to treat regular types (like $\text{Bool}$) as type constructors of arity 0 (that is, type constructors with kind $*$). In effect, they take zero arguments and produce types (themselves). ] such as $\text{List} \; \text{Int}$ or $\text{Bool}$.
 A function from one type to another, like $\text{List} \; a \rightarrow \text{Int}$.
Polytypes (type schemes) in our system can be all of the above, but may also include a “forall” quantifier at the front, generalizing the type (like $\forall a \; . \; \text{List} \; a \rightarrow \text{Int}$).
Let’s start implementing all of this. Why don’t we start with the change to the syntax of our language? We have complicated the situation quite a bit. Let’s take a look at the old grammar for data type declarations (this is going back as far as part 2). Here, $L_D$ is the nonterminal for the things that go between the curly braces of a data type declaration, $D$ is the nonterminal representing a single constructor definition, and $L_U$ is a list of zero or more uppercase variable names:
$\begin{aligned} L_D & \rightarrow D \; , \; L_D \\ L_D & \rightarrow D \\ D & \rightarrow \text{upperVar} \; L_U \\ L_U & \rightarrow \text{upperVar} \; L_U \\ L_U & \rightarrow \epsilon \end{aligned}$This grammar was actually too simple even for our monomorphically typed language! Since functions are not represented using a single uppercase variable, it wasn’t possible for us to define constructors that accept as arguments anything other than integers and userdefined data types. Now, we also need to modify this grammar to allow for constructor applications (which can be nested). To do all of these things, we will define a new nonterminal, $Y$, for types:
$\begin{aligned} Y & \rightarrow N \; ``\rightarrow" Y \\ Y & \rightarrow N \end{aligned}$We make it rightrecursive (because the $\rightarrow$ operator is rightassociative). Next, we define a nonterminal for all types except those constructed with the arrow, $N$.
$\begin{aligned} N & \rightarrow \text{upperVar} \; L_Y \\ N & \rightarrow \text{typeVar} \\ N & \rightarrow ( Y ) \end{aligned}$The first of the above rules allows a type to be a constructor applied to zero or more arguments (generated by $L_Y$). The second rule allows a type to be a placeholder type variable. Finally, the third rule allows for any type (including functions, again) to occur between parentheses. This is so that higherorder functions, like $(a \rightarrow b) \rightarrow a \rightarrow a$, can be represented.
Unfortunately, the definition of $L_Y$ is not as straightforward as we imagine. We could define
it as just a list of $Y$ nonterminals, but this would make the grammar ambigous: something
like List Maybe Int
could be interpreted as “List
, applied to types Maybe
and Int
”, and
“List
, applied to type Maybe Int
”. To avoid this, we define a “type list element” $Y’$,
which does not take arguments:
We then make $L_Y$ a list of $Y’$:
$\begin{aligned} L_Y & \rightarrow Y' \; L_Y \\ L_Y & \rightarrow \epsilon \end{aligned}$Finally, we update the rules for the data type declaration, as well as for a single constructor. In these new rules, we use $L_T$ to mean a list of type variables. The rules are as follows:
$\begin{aligned} T & \rightarrow \text{data} \; \text{upperVar} \; L_T = \{ L_D \} \\ D & \rightarrow \text{upperVar} \; L_Y \\ \end{aligned}$Those are all the changes we have to make to our grammar. Let’s now move on to implementing the corresponding data structures. We define a new family of structs, which represent types as they are received from the parser. These differ from regular types in that they do not necessarily represent valid types; validating types requires two passes, whereas parsing is done in a single pass. We can define our parsed types as follows:


We define the conversion method to_type
, which requires
a set of type variables that are allowed to occur within a parsed
type (which are the variables specified on the left of the =
in the data type declaration syntax), and the environment in which to
look up the arities of any type constructors. The implementation is as follows:


Note that this definition requires a new type
subclass, type_app
, which
represents type application. Unlike parsed_type_app
, it stores a pointer
to the type constructor being applied, rather than its name. This
helps validate the type (by making sure the parsed type’s name refers to
an existing type constructor), and lets us gather information like
which constructors the resulting type has. We define this new type as follows:


With our new data structures in hand, we can now update the grammar in our Bison file. First things first, we’ll add the type parameters to the data type definition:


Next, we add the new grammar rules we came up with:


Note in the above rules that even for typeListElement
, which
can never be applied to any arguments, we still attach a parsed_type_app
as the semantic value. This is for consistency; it’s easier to view
all types in our system as applications to zero or more arguments,
than to write coercions from nonapplied types to types applied to zero
arguments.
Finally, we define the types for these new rules at the top of the file:


This concludes our work on the parser, but opens up a whole can of worms
elsewhere. First of all, now that we introduced a new type
subclass, we must
ensure that type unification still works as intended. We therefore have
to adjust the type_mgr::unify
method:


In the above snippet, we add a new ifstatement that checks whether or
not both types being unified are type applications, and if so, unifies
their constructors and arguments. We also extend our type equality check
to ensure that both the names and arities of types match
[note:
This is actually a pretty silly measure. Consider the following three
propositions:
1) types are only declared at the toplevel scope.
2) if a type is introduced, and another type with that name already exists, we throw an error.
3) for name equality to be insufficient, we need to have two declared types
with the same name. Given these propositions, it will not be possible for us to
declare two types that would confuse the name equality check. However,
in the near future, these propositions may not all hold: if we allow
let/in
expressions to contain data type definitions,
it will be possible to declare two types with the same name and arity
(in different scopes), which would still confuse the check.
In the future, if this becomes an issue, we will likely move to unique
type identifiers.
]
Note also the more basic fact that we added arity
to our type_base
,
[note:
You may be wondering, why did we add arity to base types, rather than data types?
Although so far, our language can only create type constructors from data type definitions,
it's possible (or even likely) that we will have
polymorphic builtin types, such as
the IO monad.
To prepare for this, we will allow our base types to be type constructors too.
]
Jut as we change type_mgr::unify
, we need to change type_mgr::find_free
to include the new case of type_app
. The adjusted function looks as follows:


There is another adjustment that we have to make to our type code. Recall that we had code that implemented substitutions: replacing free variables with other types to properly implement our type schemes. There was a bug in that code, which becomes much more apparent when the substitution system is put under more pressure. Specifically, the bug was in how type variables were handled.
The old substitution code, when it found that a type
variable had been bound to another type, always moved on to perform
a substitution in that other type. This wasn’t really a problem then, since
any type variables that needed to be substituted were guaranteed to be
free (that’s why they were put into the “forall” quantifier). However, with our
new system, we are using userprovided type variables (usually a
, b
, and so on),
which have likely already been used by our compiler internally, and thus have
been bound to something. That something is irrelevant to us: when we
perform a substitution on a userdefined data type, we know that our a
is
free, and should be substitited. In short, precedence should be given to
substituting type variables, rather than resolving them to what they are bound to.
To make this adjustment possible, we need to make substitute
a method of type_manager
,
since it will now require an awareness of existing type bindings. Additionally,
this method will now perform its own type resolution, checking if a type variable
needs to be substitited between each step. The whole code is as follows:


That’s all for types. Definitions, though, need some work. First of all,
we’ve changed our parser to feed our constructor
type a vector of
parsed_type_ptr
, rather than std::string
. We therefore have to update
constructor
to receive and store this new vector:


Similarly, definition_data
itself needs to accept the list of type
variables it has:


We then look at definition_data::insert_constructors
, which converts
constructor
instances to actual constructor functions. The code,
which is getting pretty complciated, is as follows:


In the above snippet, we do the following things:
 We first create a set of type variables that can occur
in this type’s constructors (the same set that’s used
by the
to_type
method we saw earlier). While doing this, we ensure a type variable is not used twice (this is not allowed), and add each type variable to the final return type (which is something likeList a
), in the order they occur.  When the variables have been gathered into a set, we iterate
over all constructors, and convert them into types by calling
to_type
on their arguments, then assembling the resulting argument types into a function. This is not enough, however, [note: This is also not enough because without generalization using "forall", we are risking using type variables that have already been bound, or that will be bound. Even ifa
has not yet been used by the typechecker, it will be once the type manager generates its first type variable, and things will go south. If we, for some reason, wanted type constructors to be monomorphic (but generic, with type variables) we'd need to internally instnatiate fresh type variables for every userdefined type variable, and substitute them appropriately. ] as we have discussed above with $\text{Nil}$ and $\text{Cons}$. To accomodate for this, we also add all type variables to the “forall” quantifier of a new type scheme, whose monotype is our newly assembled function type. This type scheme is what we store as the type of the constructor.
This was the last major change we have to perform. The rest is cleanup: we have switched
our system to dealing with type applications (sometimes with zero arguments), and we must
bring the rest of the compiler up to speed with this change. For instance, we update
ast_int
to create a reference to an existing integer type during typechecking:


Similarly, we update our code in typecheck_program
to use
type applications in the type for binary operations:


Finally, we update ast_case
to unwrap type applications to get the needed constructor
data from type_data
. This has to be done in ast_case::typecheck
, as follows:


Additionally, a similar change needs to be made in ast_case::compile
:


That should be all! Let’s try an example:
data List a = { Nil, Cons a (List a) }
data Bool = { True, False }
defn length l = {
case l of {
Nil > { 0 }
Cons x xs > { 1 + length xs }
}
}
defn main = { length (Cons 1 (Cons 2 (Cons 3 Nil))) + length (Cons True (Cons False (Cons True Nil))) }
The output:
Result: 6
Yay! Not only were we able to define a list of any type, but our length
function correctly
determined the lengths of two lists of different types! Let’s try an example with the
classic fold
functions:
data List a = { Nil, Cons a (List a) }
defn map f l = {
case l of {
Nil > { Nil }
Cons x xs > { Cons (f x) (map f xs) }
}
}
defn foldl f b l = {
case l of {
Nil > { b }
Cons x xs > { foldl f (f b x) xs }
}
}
defn foldr f b l = {
case l of {
Nil > { b }
Cons x xs > { f x (foldr f b xs) }
}
}
defn list = { Cons 1 (Cons 2 (Cons 3 (Cons 4 Nil))) }
defn add x y = { x + y }
defn sum l = { foldr add 0 l }
defn skipAdd x y = { y + 1 }
defn length l = { foldr skipAdd 0 l }
defn main = { sum list + length list }
We expect the sum of the list [1,2,3,4]
to be 10
, and its length to be 4
, so the sum
of the two should be 14
. And indeed, our program agrees:
Result: 14
Let’s do one more example, to test types that take more than one type parameter:
data Pair a b = { MkPair a b }
defn fst p = {
case p of {
MkPair a b > { a }
}
}
defn snd p = {
case p of {
MkPair a b > { b }
}
}
defn pair = { MkPair 1 (MkPair 2 3) }
defn main = { fst pair + snd (snd pair) }
Once again, the compiled program gives the expected result:
Result: 4
This looks good! We have added support for polymorphic data types to our compiler.
We are now free to move on to let/in
expressions, lambda functions, and Input/Output,
as promised, starting with part 12  let/in
and lambdas!