The Amulet Programming Language

Posted on January 18, 2018

As you might have noticed, I like designing and implementing programming languages. This is another of these projects. Amulet is a strictly-evaluated, statically typed impure roughly functional programming language with support for parametric data types and rank-1 polymorphism à la Hindley-Milner (but no let-generalization), along with row-polymorphic records. While syntactically inspired by the ML family, it’s a disservice to those languages to group Amulet with them, mostly because of the (present) lack of modules.

Planned features (that I haven’t even started working on, as of writing this post) include generalized algebraic data types, modules and modular implicits, a reworked type inference engine based on OutsideIn(X)1 to support the other features, and, perhaps most importantly, a back-end that’s not a placeholder (i.e. something that generates either C or LLVM and can be compiled to a standalone executable).

The compiler is still very much a work in progress, and is actively being improved in several ways: Rewriting the parser for efficiency concerns (see Lexing and Parsing), improving the quality of generated code by introducing more intermediate representations, and introducing several optimisations on the one intermediate language we do have.

The Technical Bits

In this section, I’m going to describe the implementation of the compiler as it exists at the time of writing - warts and all. Unfortunately, we have a bit too much code for all of it to fit in this blag post, so I’m only going to include the horribly broken bits here, and leave the rest out. Of course, the compiler is open source, and is available on my GitHub.

Lexing and Parsing

To call what we have a lexer is a bit of an overstatement: The Parser.Lexer module, which underpins the actual parser, contains only a handful of imports and some definitions for use with Parsec’s Text.Parsec.Token module; Everything else is boilerplate, namely, declaring, at top-level, the functions generated by makeTokenParser.

Our parser is then built on top of this infrastructure (and the other combinators provided by Parsec) in a monadic style. Despite having chosen to use strict Texts, many of the Parsec combinators return Chars, and using the Alternative type class’ ability to repeat actions makes linked lists of these - the dreaded String type. Due to this, and other inefficiencies, the parser is ridiculously bad at memory management.

However, it does have some cute hacks. For example, the pattern parser has to account for being used in the parsing of both match and fun - in the former, destructuring patterns may appear without parenthesis, but in the latter, they must be properly parenthesised: since fun may have multiple patterns, it would be ambiguous if fun Foo x -> ... is destructuring a Foo or takes two arguments.

Instead of duplicating the pattern parser, one for matches and one for function arguments, we instead parametrised the parser over needing parenthesis or not by adding a rank-2 polymorphic continuation argument.

patternP :: (forall a. Parser a -> Parser a) -> Parser Pattern'
patternP cont = wildcard <|> {- some bits omitted -} try destructure where
  destructure = withPos . cont $ do
    ps <- constrName
    Destructure ps <$> optionMaybe (patternP id)

When we’re parsing a pattern match-style, the continuation given is id, and when we’re parsing an argument, the continuation is parens.

For the aforementioned efficiency concerns, however, we’ve decided to scrap the Parsec-based parser and move to an Alex/Happy based solution, which is not only going to be more maintainable and more easily hackable in the future, but will also be more efficient overall. Of course, for a toy compiler such as this one, efficiency doesn’t matter that much, but using one and a half gigabytes to compile a 20-line file is really bad.


To simplify scope handling in both the type checker and optimiser, after parsing, each variable is tagged with a globally unique integer that is enough to compare variables. This also lets us use more efficient data structures later in the compiler, such as VarSet, which stores only the integer identifier of a variable in a big-endian Patricia tree2.

Our approach, described in Secrets of the Glasgow Haskell Compiler inliner as “the Sledgehammer”, consists of duplicating every bound variable to avoid name capture problems. However, while the first of the listed disadvantages surely does apply, by doing all of the renaming in one go, we mostly avoid the latter. Of course, since then, the Haskell ecosystem has evolved significantly, and the plumbing required is a lot less intrusive.

In our compiler, we use MTL-style classes instead of concrete monad transformer stacks. We also run every phase after parsing in a single GenT monad, which provides a fresh supply of integers for names. “Plumbing” the fresh name supply, then, only involves adding a MonadGen Int m constraint to the context of functions that need it.

Since the string component of parsed names is not thrown away, we also have to make up strings themselves. This is where another cute hack comes in: We generate, lazily, an infinite stream of names that goes ["a" .. "z", "aa" .. "az", "ba" .. "bz", ..], then use the MonadGen counter as an index into that stream.

alpha :: [Text]
alpha = map T.pack $ [1..] >>= flip replicateM ['a'..'z']


The desugarer is a very simple piece of code which, through use of Scrap Your Boilerplate-style generic programming, traverses the syntax tree and rewrites nodes representing syntax sugar to their more explicit versions.

Currently, the desugarer only expands sections: That is, expressions of the form (+ e) become fun x -> x + e (where e is a fresh name), expressions like (e +) become fun x -> e + x, and expressions like .foo becomes fun x ->

This is the only component of the compiler that I can reasonably include, in its entirety, in this post.

desugarProgram = everywhereM (mkM defaults) where
  defaults :: Expr Parsed -> m (Expr Parsed)
  defaults (BothSection op an) = do
    (ap, ar) <- fresh an
    (bp, br) <- fresh an
    pure (Fun ap (Fun bp (BinOp ar op br an) an) an)
  defaults (LeftSection op vl an) = do
    (cap, ref) <- fresh an
    pure (Fun cap (BinOp ref op vl an) an)
  defaults (RightSection op vl an) = do
    (cap, ref) <- fresh an
    pure (Fun cap (BinOp vl op ref an) an)
  defaults (AccessSection key an) = do
    (cap, ref) <- fresh an
    pure (Fun cap (Access ref key an) an)
  defaults x = pure x

Type Checking

By far the most complicated stage of the compiler pipeline, our inference algorithm is modelled after Algorithm W (extended with kinds and kind inference), with constraint generation and solving being two separate steps.

We first traverse the syntax tree, in order, making up constraints and fresh type variables as needed, then invoke a unification algorithm to produce a substitution, then apply that over both the generated type (a skeleton of the actual result) and the syntax tree (which is explicitly annotated with types everywhere).

The type inference code also generates and inserts explicit type applications when instancing polymorphic types, since we internally lower Amulet into a System F core language with explicit type abstraction and application. We have TypeApp nodes in the syntax tree that never get parsed or renamed, and are generated by the type checker before lowering happens.

Our constraint solver is quite rudimentary, but it does the job nicely. We operate with a State monad with the current substitution. When we unify a variable with another type, it is added to the current substitution. Everything else is just zipping the types together. When we try to unify, say, a function type with a constructor, that’s an error. If a variable has already been added to the current substitution and encounter it again, the new type is unified with the previously recorded one.

unify :: Type Typed -> Type Typed -> SolveM ()
unify (TyVar a) b = bind a b
unify a (TyVar b) = bind b a
unify (TyArr a b) (TyArr a' b') = unify a a' *> unify b b'
unify (TyApp a b) (TyApp a' b') = unify a a' *> unify b b'
unify ta@(TyCon a) tb@(TyCon b)
  | a == b = pure ()
  | otherwise = throwError (NotEqual ta tb)

This is only an excerpt, because we have very complicated types.

Polymorphic Records

One of Amulet’s selling points (if one could call it that) is its support for row-polymorphic records. We have two types of first-class record types: closed record types (the type of literals) and open record types (the type inferred by record patterns and field getters.). Open record types have the shape { 'p | x_n : t_n ... x_n : t_n }, while closed records lack the type variable 'p.

Unification of records has 3 cases, but in all 3 cases it is checked that fields present in both records have unifiable types.

As an example, { x = 1 } has type { x : int }, the function fun x -> has type { 'p | foo : 'a } -> 'a, and (fun r -> r.x) { y = 2 } is a type error3.

No Let Generalisation

Vytiniotis, Peyton Jones and Schrijvers argue4 that HM-style let generalisation interacts badly with complex type system extensions such as GADTs and type families, and should therefore be omitted from such systems. In a deviation from the paper, GHC 7.2 reintroduces let generalisation for local definitions that meet some criteria5.

Here’s the rule. With -XMonoLocalBinds (the default), a binding without a type signature is generalised only if all its free variables are closed.

A binding is closed if and only if

We, however, have chosen to follow that paper to a tee. Despite not (yet!) having any of those fancy type system features that interact poorly with let generalisation, we do not generalise any local bindings.


After type checking is done (and, conveniently, type applications have been left in the correct places for us by the type checker), Amulet code is converted into an explicitly-typed intermediate representation, in direct style, which is used for (local) program optimisation. The AST is simplified considerably: from 19 constructors to 9.

Type inference is no longer needed: the representation of core is packed with all the information we need to check that programs are type-correct. This includes types in every binder (lambda abstractions, lets, pattern bindings in match), big-lambda abstractions around polymorphic values (a \(\lambda\) binds a value, while a \(\Lambda\) binds a type), along with the already mentioned type applications.

Here, code also gets the error branches for non-exhaustive match expressions, and, as a general rule, gets a lot uglier.

let main _ = (fun r -> r.x) { x = 2 }

(* Is elaborated into *)

let  main : ∀ 'e. 'e -> int =
  Λe : *. λk : 'e. match k {
    (p : 'e) : 'e -> (λl : { 'g | x : int }. match l {
      (r : { 'g | x : int }) : { 'g | x : int } -> match r {
        { (n : { 'g | x : int }) | x = (m : int) } : { 'g | x : int } -> m
      (o : { 'g | x : int }) : { 'g | x : int } ->
        error @int "<test>[1:15 .. 1:27]"
    }) ({ {} | x : int = 2 });
    (q : 'e) : 'e -> error @int "<test>[1:14 .. 1:38]"


As the code we initially get from lowering is ugly and inefficient - along with being full of the abstractions functional programs have by nature, it is full of redundant matches created by e.g. the fact that functions can not do pattern matching directly, and that field access gets reduced to pattern matching - the optimiser’s job is to make it prettier, and more efficient.

The optimiser works by applying, in order, a series of local transformations operating on individual sub-terms to produce an efficient program, 25 times. The idea of applying them several times is that, when a simplification pass kicks in, more simplification opportunities might arise.

dropBranches, foldExpr, dropUselessLets

These trivial passes remove similarly trivial pieces of code that only add noise to the program. dropBranches will do its best to remove redundant arms from a match expression, such as those that appear after an irrefutable pattern. foldExpr reduces uses of operators where both sides are known, e.g. 2 + 2 (replaced by the literal 5) or "foo " ^ "bar" (replaced by the literal "foo bar"). dropUselessLets removes lets that bind unused variables whose right-hand sides are pure expressions.

trivialPropag, constrPropag

The Amulet optimiser does inlining decisions in two (well, three) separate phases: One is called propagation, in which a let decides to propagate its bound values into the expression, and the other is the more traditional inlining, where variables get their values from the context.

Propagation is by far the easiest of the two: The compiler can see both the definitions and all of the use sites, and could in theory decide if propagating is beneficial or not. Right now, we propagate all literals (and records made up solely of other trivial expressions), and do a round of propagation that is best described as a rule.

let { v = C e } in ... v ...
(* becomes *)
let { v' = e } in ... C v' ...

This constructor propagation allows the match optimisations to kick in more often, and is semantics preserving.


This pass identifies match expressions where we can statically determine the expression being analysed and, therefore, decide which branch is going to be taken.

match C x with
| C e -> ... e ...
(* becomes *)
... x ...


It is always safe to turn a match where the term being matched is a diverging expression into only that diverging expression, thus reducing code size several times.

match (error @int "message") with ...
(* becomes *)
error @int "message"

As a special case, when one of the arms is itself a diverging expression, we use the type mentioned in that application to error to fix up the type of the value being scrutinized.

match (error @foo "message") with
| _ -> error @bar "message 2"
(* becomes *)
error @bar "message"


This transformation turns match expressions where the expression being dissected is itself another match “inside-out”: we push the branches of the outer match “into” the inner match (what used to be the expression being scrutinized). In doing so, sometimes, new opportunities for match-of-known-constructor arise, and the code ends up simpler.

match (match x with
        | A -> B
        | C -> D) with
  | B -> e
  | D -> f
(* becomes *)
match x with
  | A -> match B with
    | B -> e
    | D -> f
  | C -> match D with
    | B -> e
    | D -> f

A clear area of improvement here is extracting the outer branches into local let-bound lambda abstractions to avoid an explosion in code size.

inlineVariable, betaReduce

In this pass, use of a variable is replaced with the definition of that variable, if it meets the following conditions:

In doing this, however, we end up with pathological terms of the form (fun x -> e) y. The betaReduce pass turns this into let x = y in e. We generate let bindings instead of substituting the variable with the parameter to maintain the same evaluation order and observable effects of the original code. This does mean that, often, propagation kicks in and gives rise to new simplification opportunities.


I was planning to write a section with a formalisation of the language’s semantics and type system, but it turns out I’m no mathematician, no matter how hard I pretend. Maybe in the future.

Our code generator is wholly uninteresting, and, most of all, a placeholder: This is why it is not described in detail (that is, at all) in this post. I plan to write a follow-up when we actually finish the native code generator.

As previously mentioned, the compiler is open source: the code is here. I recommend using the Nix package manager to acquire the Haskell dependencies, but Cabal should work too. Current work in rewriting the parser is happening in the feature/alex-happy branch.

  1. Dimitrios Vytiniotis, Simon Peyton Jones, Tom Schrijvers, and Martin Sulzmann. 2011. OutsideIn(X): Modular Type Inference With Local Assumptions. Note that, although the paper has been published in the Journal of Functional Programming, the version linked to here is a preprint.

  2. This sounds fancy, but in practice, it boils down to using Data.IntSet instead of Data.Set.

  3. As shown here. Yes, the error messages need improvement.

  4. Dimitrios Vytiniotis, Simon Peyton Jones, Tom Schrijvers. 2010. Let Should not be Generalised.

  5. As explained in this blog post.