(I've been editing this post on and off for almost a year. I'm not really happy with it, but I suspect I never will be.)
Several months ago I gave a talk at work about Hindley-Milner type inference. When I agreed to give the talk I didn't know much about the subject, so I learned about it. And now I'm writing about it, based on the contents of my talk but more fleshed out and hopefully better explained.
I call this a reckless introduction, because my main source is wikipedia. A bunch of people on the internet have collectively attempted to synthesise a technical subject. I've read their synthesis, and now I'm trying to re-synthesise it, without particularly putting in the effort to check my own understanding. I'm not going to argue that this is a good idea. Let's just roll with it.
I'm also trying to tie in some quasi-philosophy that surely isn't original to me but I don't know if or where I've encountered it before.^{1}
When people write software, sometimes it doesn't do exactly what we want. One way to find out is to try running it and see, but that's not ideal because any complicated program will have way too many possible inputs to test. (Especially when you consider that inputs include things like "amount of free space on disk" and "time taken for a web server to respond to a request".) So it would be nice if we could mathematically prove whether our software does what we want, without actually running it. Can we do that?
That's not a very well-defined question, but we can ask more precise versions of it. Here's a well-known one: given some possible input to our software, we might want to prove that our software will eventually stop running. Can we prove that?
That question is known as the halting problem, and the simple answer is that we can't, not in general; the halting problem is undecideable. But the full answer is more complicated.
To solve the halting problem, we want a program that, when shown another program and some input to be fed to that program, satisfies three different conditions:
And that's not possible. But we can compromise on any of the three. We can make a program that sometimes doesn't return an answer, or one that sometimes gets the answer wrong. But perhaps most interestingly, we can make a program that sometimes says "I don't know".
And when you allow that answer, you can create a language on which the halting problem is decideable. You can write a program that will tell you truthfully whether any program written in that language will terminate; and for any other program, will say "I don't know". (Perhaps expressed in words like "syntax error on line 1".)
Now, the halting problem is tricky. It turns out that if you create a language like that, there are a lot of interesting things that programs written in that language just won't be able to do; the language will necessarily be Turing incomplete.^{2} But there are also lots of interesting things that they can do. To give three examples of such languages^{3}:
So although these languages can't do everything, they can still be incredibly useful in their domains. More useful than a more general purpose language might be. One reason for this is that being able to prove non-termination is a useful property of the language. If you had to write a SQL query in C, it would be all too easy to write some C code that would accidentally loop forever.
I'm trying to illustrate here something that seems to me important, which is that there's a tradeoff between what I'll call expressiveness and legibility. A programming language is expressive if you can easily write many interesting programs in it^{5}; it's legible if you can easily say many interesting things about the programs you've written in it. And I claim that the most expressive programming languages won't be the most legible, and vice-versa; though there will certainly be languages which are neither expressive nor legible. This tradeoff seems fundamental to me, and I expect that some approximation of it has been proven as a theorem.^{6}
I haven't defined these very well, but hopefully some examples will help. I will also clarify that both of them are highly dimensional; and that "raw computational power" is one of the things that expressiveness can point at, but not the only thing; and "human readability" is not really one of the things that legibility points at, but many things that increase legibility will also increase human readability.
Perl-compatible regular expressions can classify sets of strings that normal regular expressions can't. But they're harder to make time and space guarantees about. And it's possible to prove whether two regular expressions are equivalent, but that's not possible in general for PCREs (proof: PCREs can encode CFGs; CFGs can't be proved equivalent).
Under certain assumptions, Haskell's monadic IO lets you look at the type of a piece of code and know that it won't depend on external state. In return, a function can only bring in external state if its caller allows it to (which requires having permission from its own caller, and so on).
The assumptions in question are false (partly because unsafePerformIO
exists), but I've been able to get away with pretending they're true (partly because unsafePerformIO
is punishable with excommunication).
Custom operators (at least as implemented in Haskell and Elm) are equivalent to named functions, and don't gain or cost much in terms of legibility and expressivity. They simply make code more or less readable. But operator overloading, at least when combined with dynamic typing, gains expressivity at the cost of legibility (you no longer know that a + b
will do anything remotely like an addition).
Macros make it easier to do things like create DSLs, reduce boilerplate, and set compile-time config options. But they mean that a function call might not look like one, or vice-versa; expressions might get evaluated many times, or not at all; and the code might perform differently depending on the phase of the moon when it was compiled.
So we've got this tradeoff, and in our programming language design we try to navigate it. We try to find kinds of legibility that can be bought for little cost in expressiveness. Or more precisely, we try to find kinds of legibility that we care about, and that can be bought for little cost in kinds of expressiveness that we care about.
And Hindley-Milner type systems are a tradeoff that's proved fairly successful, both in direct use and as inspiration. At my company^{7}, we use Elm^{8}, which runs on an approximately HM type system. (I don't think it's pure HM, due to extensible record types.) We also use Haskell^{9}, which runs on a type system that extends HM in many directions. Haskell's system is more expressive and less legible, but still successful. (I'll mostly be using Elm for examples in this post, and not extensible records.) ML and OCaml are other notable languages based on HM, though I haven't used either.
The legibility HM offers is, roughly, the ability to prove that a program typechecks. I'm not going to clarify exactly what that means, but we probably all have a decent idea. It's the thing that lets the Elm compiler say "no, that program is trying to add a string to an int, bad program", while the Python interpreter doesn't know that's going to happen until it's too late. The Elm compiler will refuse to compile your program unless it can logically prove that it will typecheck.
More precisely, what HM offers isn't type checking but the more general type inference. (And beyond that, type inference in roughly linear time.) Type inference doesn't just tell you whether a program typechecks, but what its type is; a program fails to typecheck iff no type can be inferred for it.
What this means is that there's no need to supply type annotations. And indeed, in Elm you can get away without them, except I think for extensible records. In Haskell you sometimes can't, because Haskell loses some of the legibility that HM offers.
(We typically do supply type annotations, but that's because they're useful. Partly as documentation for humans, partly to help pinpoint errors when our programs fail to typecheck.)
And so in an HM system you get no runtime type errors. And although not all runtime errors are type errors, in many cases they could be. For example, an array out-of-bounds exception isn't a type error. But when designing a language, you can decide that array out-of-bounds exceptions won't exist, any array lookup will return either a value from the array or null
. If type errors are possible, you've just eliminated one source of errors by pushing them somewhere else, and possibly somewhere harder to debug. But in HM, you've eliminated one source of errors by pushing them somewhere more visible, where they can be ruthlessly executed.
Elm actually tries to promise no runtime errors, period, provided you stay inside Elm. On one level, I think that's a fairly minor imposition on language design, something you get "for free" by deciding that none of the built-in functions you provide will ever throw a runtime error. On another level, it seems completely impractical to decide for example that cons
will return a meaningful value if it can't allocate more memory. I'm not aware that Elm even tries to handle those errors.
(Haskell doesn't try to promise the same thing, and allows functions to return undefined
. This is another legibility-expressiveness tradeoff.)
So HM's legibility gain is: type inference, powerful type system, no runtime type errors, optionally no runtime errors at all. It's good.
Meanwhile, the expressiveness cost is that you need to write your programs in ways that the type inference algorithms can work with, which forbids some things that you might like to do.
For example, suppose you want to clamp a number to between -1 and +1. In Python, you could write that like
def clamp(x): sorted([-1, x, 1])[1]
and as long as sorted
always returns a list of the same length it started with, that works fine^{10}. But it only works because the Python interpreter allows you to be reckless with array indexing. Elm doesn't let you be reckless, and so Elm has no equivalent way to perform array lookup. If you tried to write the same function in the same way in Elm, the result in the compiler's eyes would not be a number but a Maybe
number - AKA "either a number or Nothing
". (Nothing
is roughly equivalent to None
in python or null
in many other languages, but you have to explicitly flag when it's allowed.) When you actually run this code, you will always get a number and never Nothing
. But the compiler can't prove that.
(Again, I stress that you will never get Nothing
as long as your sort function always returns a list of the same length it started with. That's something you can prove for yourself, but it's not something the Elm compiler can prove. It's not even the sort of thing the Elm compiler knows can be proven. And so in turn, it can't prove that you'll never get a Nothing
here.)
And then the Elm compiler would force you to account for the possibility of Nothing
, even though there's no way that possibility could occur at runtime. One option is to pick an arbitrary result that will never be exposed. That works fine until the code goes through several layers of changes, an assumption that used to be true is now violated, and suddenly that arbitrary result is wreaking havoc elsewhere. Or in Haskell, your program is crashing at runtime.
To be clear, that's not significantly worse than what we get in Python, where the code can also go through several layers of changes that result in it crashing at runtime. But we were hoping for better.
And in this case "better" is easy enough, you can just write your function to avoid indexing into a list, and then it can return a number with no need for trickery. The point isn't that you can't do the thing. The point is that (a), even if the thing is safe, the compiler might not know that; (b), if you decide it's safe anyway and find some way to trick the compiler, the compiler no longer protects you; and (c), if you want to do it in a way the compiler knows is safe, you might need to put in some extra work.
For another example, HM type systems can't implement heterogenous lists. So this is really easy in python:
def stringify_list(l):
return [ repr(x) for x in l ]
stringify_list(["hello",
0,
["here's a", "nested list", {"and": "maybe", "a": "dict"}],
"it can even be passed itself, like so:",
stringify_list])
but impossible in Elm. You can sort of get the same effect by creating a type with many constructors
type HeteroType = HTInt Int
| HTString String
| HTBool Bool
| HTList (List HeteroType)
| ...
but it's not quite the same, because it can only accept types you know about in advance. Also, it's a massive pain to work with.
For a third example: Haskell is known for its monads. But Elm has no equivalent, because an HM type system can't support generic monad programming. You can implement the generic monad functions for specific cases, so there's Maybe.map
and List.map
, but there's no equivalent of Haskell's fmap
which works on all monads.
I've talked about the tradeoffs that HM type systems offer, but not what HM type systems actually are. So here is where I get particularly reckless.
This bit is more formal than the rest. It's based on the treatment at wikipedia, but I've tried to simplify the notation. I'm aiming for something that I would have found fairly readable several months ago, but I no longer have access to that version of me.
Also, this part is likely to make more sense if you're familiar with at least one HM-based language. That's not a design feature, I just don't trust myself to bridge that inferential gap.
For an HM system, you need a language to run type inference on, and you need types to run type inference with, and you need some way to combine the two. You could use the language with no type inference, if you didn't mind crashes or weird behaviour at runtime, when you made a mistake with typing. (Haskell allows this with a compiler option.^{11}) And you could run type inference without caring about the semantics of the language, treating it as essentially a SuDoku, an interesting puzzle but meaningless. (Haskell supports this, too.) But by combining them, the semantics of the language are constrained by the type system, and runtime type errors are eliminated.
Types come in a conceptual hierarchy which starts with type constants. That's things like, in Elm, Int
, Float
, Bool
, String
, Date
, ()
. It also includes type variables, which in Elm are notated with initial lower-case, like a
and msg
. (Though the type variables number
, comparable
and appendable
are special cases that I won't cover here.)
Next in the type hierarchy is applied types. Here a "type function" is applied to arguments, which are type constants and/or other applied types. These are things like List Int
, Maybe (List Float)
, Result () Date
, and a -> String
. (In that last one, the type function is the arrow; Haskell would allow you to write it (->) a String
. Aside, (->)
is the only type that HM specifically requires to exist; a -> b
is the type of functions taking a parameter of type a
and evaluating to a result of type b
.) Notably, an applied type must have a specific named type function as its root; you can't have m Int
, which you would need for generalised monads.
Type constants and applied types are monotypes. You get a polytype by optionally sticking one or more "∀"s in front of a monotype. ("∀" is pronounced "for all", and in Haskell can be written forall
.) So for example a -> Int
is a monotype, but ∀a. a -> Int
is a polytype. So is ∀a. ∀b. a -> Int -> b
, which is written equivalently as ∀a b. a -> Int -> b
. ∀b. a -> Int
is also a polytype; since the quantified variable doesn't show up, it's equivalent to the monotype a -> Int
. We can do something like that to any monotype, so for simplicity we might as well decide that monotypes count as a special case of polytypes, not as a distinct set.
Type signatures in Elm typically have an implied "∀" over whichever variables it makes sense to quantify. (There's no syntax for explicitly writing the "∀".) So the type of List.map
would be written
map : (a -> b) -> List a -> List b
but I'll be writing
map : ∀a b. (a -> b) -> List a -> List b
for clarity. Because there's one place where Elm doesn't give an implied ∀, which is when you have scoped types. To demonstrate by example,
const : ∀a b. a -> b -> a
const x = let foo : b -> a
foo y = x
in foo
const
has a polytype here, but foo
has a monotype, because (in context) its argument type and return type are constrained. If you tried to swap a
and b
in the type signature for foo
, or rename either of them, the Elm compiler would complain.
The language has four kinds of expression, and each has a rule relating it to the type system. You need variables and constants, function calls, lambda expressions, and let statements.
Variables and constants are things like True
, 0.2
, Just
, "Hello"
, []
, ()
, List.map
. Each of these has a declared type, which in Elm is notated with :
. So True : Bool
, 0.2 : Float
, Just : ∀a. a -> Maybe a
, "Hello": String
, [] : ∀a. List a
, () : ()
, List.map : ∀a b. (a -> b) -> List a -> List b
.
The rule that relates these to the type system is that type declarations imply type judgments. Mathematically it looks like
$$ \frac{x : π \quad π ⊑ μ}{x \sim μ}. $$
Reading clockwise from top left, this says: if you have a variable $x$ declared to have some polytype $π$, and if the monotype $μ$ is a specialisation of $π$, then $x$ can be judged to have type $μ$. ($π$ always denotes a polytype, and $μ$ always denotes a monotype.)
A type judgment, as opposed to a declaration, provides a type that an expression can be used as. A judgment is always as a monotype.
And type specialisation, denoted $⊑$, is the process of replacing quantified variables with less-quantified ones. So for example the type ∀a b. a -> b -> a
might be specialized to ∀a. a -> String -> a
, or to ∀b. Int -> b -> Int
; and from either of those, it could be further specialised to Int -> String -> Int
. Of course String -> Int -> String
and List Float -> (Float -> String) -> List Float
are valid specialisations too.
Thus: we have the type declaration [] : ∀a. List a
, and we have (∀a. List a) ⊑ List Int
, and so we can form the type judgment [] ~ List Int
. We also have (∀a. List a) ⊑ List String
, and so [] ~ List String
. And [] ~ List (List (Maybe Bool))
, and so on.
Function calls are things like not True
, (+ 1)
, List.Map Just
. And the rule relating them to the type system is that function calls consume function types. This is the simplest of the rules. Mathematically it looks like
$$ \frac{f \sim μ → μ' \quad v \sim μ}{f v \sim μ'}. $$
Or: if $f$ can be judged to have a function type $μ → μ'$, and $v$ can be judged to have type $μ$, then the function call $fv$ can be judged to have type $μ'$.
Thus: we can infer the type judgment toString ~ (Int -> String)
, and we can infer 3 ~ Int
, and so we can infer toString 3 ~ String
.
Also, we can infer List.map ~ ((Int -> Maybe Int) -> (List Int -> List (Maybe Int)))
, and we can infer Just ~ (Int -> Maybe Int)
. So we can infer List.map Just ~ (List Int -> List (Maybe Int))
Lambda expressions are things like \x -> Just x
, and in Elm they're used implicitly when something like const x y = x
is turned into const = \x -> \y -> x
. The type system rule is that lambda expressions produce function types. Mathematically:
$$ \frac{x : μ ⇒ e \sim μ'}{λx.e \sim μ → μ'}. $$
Or: suppose that the type declaration $x : μ$ would allow us to infer the judgment $e \sim μ'$. In that case, we could judge that $λx.e \sim (μ → μ)'$.
Typically $e$ would be some expression mentioning the variable $x$, but it's no problem if not. In that case, if you can get $e \sim μ'$ at all, you can get it assuming any $x : μ$, and so you have $λx.e \sim (\mathtt{Int} → μ')$ and $λx.e \sim (\mathtt{String} → μ')$ and $λx.e \sim (\mathtt{Result\ String\ (List\ (Maybe\ Float))} → μ')$ and so on.
Thus: given the declaration x : Int
, we can infer the judgment [x] ~ List Int
. And so we can infer the judgment (\x -> [x]) ~ (Int -> List Int)
.
Let expressions read like let x = y in a
. Semantically, this is very similar to using a lambda expression, (\x -> a) y
. But HM treats them differently in the type system, allowing a let expression to introduce polytypes. That permits code like
let f x = [x]
in (f "", f True)
-- returns ([""], [True])
If you tried to rewrite this as a lambda, you would get
(\f -> (f "", f True))(\x -> [x])
But type inference fails here, because there's no monotype declaration for f
that allows a type judgment for (f "", f True)
. So the precondition for the lambda rule never obtains, and so in turn, no type judgment can be made for the expression \f -> (f "", f True)
.
Let expressions compensate for this deficiency, with the rule let expressions are like polymorphic lambda applications. (I don't have a good name for it.) Mathematically:
$$ \frac{a \sim μ \quad x : \bar{μ} ⇒ b \sim μ'}
{(\mathtt{let}\ x = a\ \mathtt{in}\ b) \sim μ'} $$
Or: suppose that $a$ can be judged to have type $μ$, and that the declaration $x : \bar{μ}$ would allow us to infer the judgment $b \sim μ'$. In that case, we could judge that $(\mathtt{let}\ x = a\ \mathtt{in}\ b)$ has type $μ'$.
This introduces the notation $\bar{μ}$, which generalises a monotype to a polytype. How it works is: if $μ$ mentions a type variable $a$, and $a$ isn't quantified over in the surrounding context, then $\bar{μ}$ contains a "$∀a$".
Thus: we can infer (\x -> [x]) ~ (a -> List a)
, where a
is a type variable unused in the surrounding context. That type generalises to ∀a. a -> List a
. And given the declaration f : ∀a. a -> List a
, we can infer (f "", f True) ~ (List String, List Bool)
. So in total, we can infer
$$ (\mathtt{let\ f\ x\ =\ [x]\ in\ (f\ "",\ f\ True)})
\sim \mathtt{(List\ String,\ List\ Bool)}. $$
(It seems a little strange to me that the approach here is to first construct a meaningless type, and then quantify over it. Still, that's my understanding. It's of course possible I'm mistaken.)
Why do we need both let
and lambda? Well, we can't replace lambda expressions with let expressions: they're not re-usable. (When you translate a let expression into a lambda expression, you actually generate a lambda applied to an argument. There's no way to translate a lambda expression by itself into a let expression.) Meanwhile, I'm not entirely sure why we can't make lambdas polymorphic in the same way let expressions are. I think the answer is that if we tried it, we'd lose some of the legibility that HM offers - so let can be more powerful in the type system because it's less powerful in the language. But I'm not sure exactly what legibility would be lost.
There's an interesting thing about the system I just described: it may or may not be Turing complete.
The problem is that there's no specified way of doing recursion. A function can't call itself, and it can't call any other function that can call it.
But a fixed-point combinator allows recursion, and might be included in the initial set of variables. Failing that, the proper recursive types can be used to define one. (Elm and Haskell allow us to define such types^{12}.)
Failing both of those, we can introduce a new kind of expression
$$ \frac{x : μ ⇒ a \sim μ \quad x : \bar{μ} ⇒ b \sim μ'}
{(\mathtt{letrec}\ x = a\ \mathtt{in}\ b) \sim μ'}. $$
This is much the same as let
, but makes the variable x = a
available when evaluating a
. It's only available as a monotype when evaluating a
, and still doesn't get generalised to a polytype until evaluating b
.
(Elm and Haskell provide letrec
as let
and don't provide simple let
at all.)
But if an HM language doesn't provide the appropriate variables or types, and doesn't implement letrec
or something similar, it won't be Turing complete. Legibility gain, expressivity cost.
And modulo some small details, that's the entirety of a Hindley-Milner type system. If you have a language with those features, and a suitable set of types, you can perform type inference.
What we have is a set of rules that allows us to construct proofs. That is, if we look at a program written in this language, we would be able to construct a proof of its type (or lack thereof). But I already said HM is better than that: it lets us mechanically construct a proof, in (roughly) linear time.
I confess, I'm not entirely sure how to do that. The outline is obvious, recurse down the parse tree and at each step apply the appropriate rule. But since a constant can be judged as one of many types, you need to keep track of which types are acceptable. Wikipedia hints at how it works, but not in a way that I understand particularly well.
Elm and Haskell both support many things not covered so far. To look at some of them briefly, and occasionally getting even more recklesss,
It seems obvious, but both allow you to evaluate the language, something I haven't touched on much. And it does need to be touched on, because there's more than one way to do it. Haskell uses a lazy evaluation model, while Elm is strict.
Both have ways to introduce new types. That doesn't change what we've seen, but it does separate the languages into two parts. One part describes the types used in a program and one part implements the semantics of a program.
Both also support case statements along with destructuring, like
mHead : Maybe (List a) -> Result Bool a
mHead ml = case ml of
Just (a::_) -> Ok a
Just _ -> Err True
Nothing -> Err False
To implement these, you'd want to add a fifth class of language expression. But I think it would be possible in theory to write a "thin" first-pass compiler to translate these statements into the existing language. By "thin" I mean to do this in such a way that we don't lose any of the legibility guarantees we care about.^{13} (For example, if this compiler turned $n$ bytes of code in a case statement into more than $O(n)$ bytes of code in the base language, or if it ran in more than O(n) time, this condition would fail.)
If I'm right about that, then case statements neither make the language more expressive nor less legible, at least in one important sense.
(By comparison, if-then-else statements are also another class of language expression, but one which can obviously be thinly compiled down to the existing ones.)
In the type system, Elm supports record types, which are a lot like tuples but with nicer syntax. I believe these too could be thinly compiled down. But it also supports extensible records, which are more complicated. On one level you can think of a type like {a | x : Int, y : Int}
like a tuple ∀a. (a, Int, Int)
. But then this tuple needs to be unpacked and manipulated when you pass it into a function expecting an {a | x : Int}
.
I believe this is unresolvable, and extensible records represent an extension of Elm from HM. (But one with fairly minor legibility costs, in comparison to the expressiveness gains.)
Haskell supports typeclasses, which are a way of allowing functions to operate on multiple different types. (For example, the show
function can be applied to a String
, an Int
, a ()
, a [Float]
, ….) Elm doesn't, but simple typeclasses can be emulated with only a little added verbosity.
Another thing I'll say is that I've been talking about legibility and expressivity of a language. But a type system is itself a language, and may be more or less legible and expressive. I don't have a strong intuition for how these interact.
There's a lot more I could add to this post. Some things that I omitted for brevity, some that I omitted because I don't know enough about them yet^{14}, and some that I omitted because I don't know about them at all. I don't know what a sensible cutoff point is, so I'm just going to end it here.
From writing my original talk, and subsequently this blog post, I think I understand HM type systems much better than I used to. Hopefully you think the same. Hopefully we're both correct. If you see any inaccuracies, please point them out.
While writing this essay I came across the talk Constraints Liberate, Liberties Constrain. From the title and the context I encountered it, it sounds like it's on the same subject. But I haven't watched it, because it's in the form of a 50 minute video. ↩
If the halting problem is decideable on a language, the language is Turing incomplete. I don't know whether the reverse is true: are there Turing incomplete languages on which the halting problem is still undecideable? I'm mostly going to assume not. At any rate, I don't think I'm going to discuss any such languages. ↩
To nitpick myself: these aren't just languages for which you can prove termination, they're languages which never terminate, at least not for finite inputs. I don't offhand know any languages which are Turing incomplete but have the ability to loop forever, though such a thing can exist. ↩
Specifically, it looks to me like SQL-99 without recursive common table expressions is Turing incomplete. I've only ever used nonrecursive CTEs. ↩
I've subsequently discovered that wikipedia uses the same name for this concept. ↩
I think this is related to the way that ZF set theory can encode Peano arithmetic. Thus, ZF is more expressive than PA. But because ZF allows you to construct objects that PA doesn't, there are more things you can say about "all objects in PA" than about "all objects in ZF". So PA is more legible than ZF. I don't understand the Curry-Howard correspondence, but I think that's related too. ↩
"My company" is a phrase which sometimes means "the company I own or run" and sometimes "the company I work for". Here it means the latter. I don't know an unambigous way to phrase that which I don't find slightly awkward, so instead I'm using a super-awkward footnote. But, y'know. Ironically, or something. ↩
We use Elm 0.18. 0.19 is a fairly significant version change, but I think not different enough to be relevant for this post. ↩
Specifically GHC, which offers many extensions over Haskell proper. Whenever I refer to Haskell, I'm really talking about the language that GHC implements. ↩
At any rate, it works fine when you pass it a number. If you pass it something else, it might do anything. ↩
Well, sort of. It still performs type inference, it just allows it to fail. I'm not sure if "no type inference at all" would work for Haskell; but I do think it would work for a pure HM system, if you count things like "3
is of type Int
" as a raw fact, non-inferred. ↩
Minor brag: I myself contributed the Elm implementation on that page. ↩
I think it might look something like this:
mHead ml =
if *isJust ml && (*fromJust ml (\_x -> *isCons _x)) then
*fromJust ml (\_x -> *fromCons _x (\a _ -> Ok a))
else if *isJust ml then
*fromJust ml (\_ -> Err True)
else if *isNothing ml then
Err False
else
*fail
functions marked with a *
can be hidden from the language user. Additionally, *fromJust
, *fromCons
and *fail
would be able to throw runtime errors. These don't violate Elm's "no runtime errors" policy, because the compiler would only generate them in contexts where it could prove they wouldn't throw. (In the case of *fail
, when it could prove that code branch was unreachable, so it could also just not bother.)
I'm very much spitballing here. I wouldn't be surprised to discover that the approach I've described is completely unworkable. ↩
Not that that stopped me from writing this entire post. ↩
Posted on 05 May 2019
Comments elsewhere: /r/haskell; LessWrong
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