int. A type system is sound if it correctly approximates the dynamic behavior and predicts its outcome: if the static semantics predicts that a term has the type
int, the dynamic evaluation of the term, if it terminates, will yield an integer.
Conventional type-checking is the big-step evaluation in the
abstract: to find a type of an expression, we fully `evaluate' its
immediate subterms to their types. We propose a different kind of type
checking that is small-step evaluation in the abstract: it unzips (pun
intended: cf. Huet's zipper) an expression into a context and a
redex. The type-checking algorithms in the paper are implemented in
Twelf; the complete code is available. In particular, the
lfix-calc.elf implements, by way of
introduction and comparison, the small-step type-checking for simply typed
The benefit of the small-step type checking is that, by `following' the control flow, it is more suitable for effectful computations. A more interesting result comes from the fact that static evaluation, unlike the dynamic one, goes `under lambda'. Thus the small-step static evaluation context is a binding context.
Colon becomes the new turnstile! Since the evaluation context is neither commutative or associative but is subject to focusing rules, we obtain the framework for expressing various substructural and modal logics.
Joint work with Chung-chieh Shan.
Commented Twelf code with small small-step type checking algorithms and many sample derivations.
Abstract interpretation in the denotational context:
Patrick Cousot: Types as abstract interpretations. POPL 1997, 316-331.
Abstract interpretation as an aid during program development
of logical programs:
Manuel Hermenegildo, Germa'n Puebla, and Francisco Bueno: Using global analysis, partial specifications, and an extensible assertion language for program validation and debugging.
In: The logic programming paradigm: A 25-year perspective, ed. Krzysztof R. Apt, Victor W. Marek, Mirek Truszczynski, and David S. Warren, 161-192, 1999. Berlin: Springer-Verlag.
Interpreting types as abstract values: A tutorial on Hindley-Milner type inference
Session types are a form of typestate: they describe not only what is communicated now but also what should or could be communicated next. Therefore, they are non-trivial to realize in conventional languages. We need not just monads but parameterized monads as well as type functions and linear types. Realizing session types in OCaml requires inordinary cleverness -- meaning the implementation and the error messages are very hard to understand. One is constantly reminded of template metaprogramming in C++.
We present the ongoing work exploring a very different approach to session typing. The user describes communication in a DSL embedded in OCaml, and writes computational parts also in OCaml. The DSL embedding again uses OCaml, rather than type system magic, to implement session type checking and inference. We still obtain static assurances about communicating programs. We can emit detailed error messages customized at will, and use the ordinary debugger to debug type errors. The key idea is staging: run-time checks in the generator are `type-checks' from the point of view of the generated program. Fancy types are replaced with boring generator terms. (In type-checking, like in investing, exciting is ruinous.) Our library of binary session types supports multiple communication channels, internal and external choices, and channel delegation.
Joint work with Keigo Imai of Gifu University.
Poster at the Programming and Programming Languages workshop PPL 2019. March 6-8, 2019. Shidotaira onsen (Hanamaki, Iwate, Japan).
fork(2), splitting the current computational process into several, one per choice alternative. Each process has its own memory, whose contents is inherited from the parent but can be mutated without affecting the other processes -- thus ensuring that independent non-deterministic choices are indeed evaluated independently. As in Unix, making copies of process memory is best done lazily, as `copy-on-write'. To manage process memory and copy-on-write, the Unix kernel maintains a special data structure. First-class memory is like such data structure, letting an ordinary programmer do the kernel-like memory management.
First-class memory has at least the following interface (described below in Haskell):
type Memory -- the abstract type of first-class memory empty :: Memory -- allocate new, initially empty memory size :: Memory -> Int -- get the current memory size (the number of allocated cells) type Ref a -- the abstract reference to a memory cell holding a value of type a new :: a -> Memory -> (Ref a, Memory) get :: Ref a -> Memory -> a set :: Ref a -> a -> Memory -> MemoryThe interface is similar to the familiar
writeSTRefoperations. However, our
setexplicitly use the
Memoryin which the cell is (to be) located.
The interface comes with the performance requirement: the execution
set (or, more precisely, the number of accesses to
the real computer memory) should not exceed a small constant. In
other words, the implementation of
Memory should at least be as
efficient as the
The challenge is implementing the interface efficiently -- and, at the same time, statically guaranteeing some degree of correctness. The correct implementation should satisfy the familiar laws:
let (r,m') = new x m in get r m' === x for all x. let m' = set r' x m in get r m' === get r m if r /= r', r' exists let m' = set r x m in get r m' === x if r exists let m' = set r x m in set r y m' === set r y mThat is: we get what we store; storing in a memory cell overrides its previous value without affecting the other cells. Ensuring these laws in types seems to require too complicated type system to be practical. Still, we would like the types at the very least guarantee an approximation of these laws, where
===is understood as equating all values of the same type. In other words, we want to ensure that
set-ting of existing references are total and type-preserving. We want to use types so that the correctness guarantees are clear even to the compiler, which would then optimize out redundant checks. Implementing practical first-class memory with such typing guarantees is still an open challenge.
Let's consider a straightforward non-solution. Recall,
Memory should store values of arbitrary types. The first impulse
is to represent memory cells as
Dynamic, `dynamically-typed' values.
Hence the implementation
type Memory = IntMap Int Dynamic type Ref a = Inta finite map from locations (
Int) to values of some type. Alas, the operation
fromDynamic :: Dynamic -> Maybe ato read from this
Dynamicreference cell is partial.
Dynamicstores a value and the representation of its type;
fromDynamicchecks that the (representation of the) requested result type
amatches the type (representation) of the stored value. Trying to extract from the
Dynamica value of a different type than was stored returns
Nothing. In the correctly implemented
getfor an existing reference is total and type-preserving. Therefore, the type-matching check of
fromDynamicshould always succeed and hence redundant. The
Dynamicrepresentation is unsatisfactory because it imposes a run-time check that should never fail. Again, the challenge is to ensure correctness statically and avoid run-time assertions.
Besides the performance,
Dynamic is unsatisfactory theoretically. A
Dynamic stores a value and a representation of its type. On the
other hand, the memory address -- the location of a memory cell --
uniquely identifying a cell also uniquely describes the type of the
stored value. That is, within given
Memory, the location is itself
a type representation: location equality implies type
equality. Therefore, the type representation used in
The pragmatic way to avoid redundant checks is to use
unsafeCoerce. Although helping performance, it forces us to rely on
pen-and-paper correctness proofs. We better make no mistakes in the
proof, or else risk a segmentation fault.
Let's recall the outline of the pen-and-paper correctness argument,
to see why it is so hard to express in types. First, the type of reference
Ref a includes the type of the stored value. The type signatures
set show that
once a reference
r is created to hold the value of some type
get r and
set r can only read and write values of that type
To complete the proof we have to check that the reference
new x m
is `fresh' -- it names a cell that
is different from any cell in the parent memory
get r m' and
set r y m' really access that very
r. The correctness of
set is hence tied up
to the third function,
new really creates fresh references.
Expressing and using this freshness property in types is the real challenge.
First-class memory module used in the probabilistic programming language Hakaru10
In this implementation the reference Ref acts as a cache for the stored value.
Probabilistic programming using first-class stores and first-class continuations
`Dead code' generally means a subexpression that has no information flow -- neither data nor control -- into the expression's result. In this article we use the term specifically for irrelevant definitions: well-typed and side-effect-free definitions with no information flow into the expression of interest; the defined identifiers do not or even cannot appear in the expression in question.
We start with the simplest, well-known example of an irrelevant definition making an expression untypeable. One has certainly seen such examples, in any language with the Hindley-Milner--like inference. For concreteness, we use the simple subset of OCaml below:
let exp = let id = fun u -> u in let f = id id in let z = fun () -> f true in f 1The code does not type-check, with the type error reported for
f 1. However, if we remove the definition of
z-- which clearly does not appear within
f 1-- the latter expression becomes well-typed.
Exercise: what other valid (from the dynamic-semantics point of view)
exp will make it well-typed?
In Haskell98 and above, the similar problem is usually associated with the infamous MonomorphismRestriction:
f = print exp = f 1 z = f TrueThe above code fails to type-check, with the compiler blaming
f 1. The definition of
zis obviously not relevant to
exp(in realistic programs, it could have appeared many lines down the program). Yet if we remove that definition, the program type-checks and runs. Although the MonomorphismRestriction is often the culprit, it is not always to blame. For example:
exp = do f <- return id r <- return (f 1) z <- return (f True) print rAlthough
zis bound after
r, it nevertheless manages to get the type-checker to complain about the
rdefinition, specifically, about the literal
1. This time, the problem has nothing to do with the MonomorphismRestriction, or the presence or absence of explicit type signatures.
In all these examples, the type of
f cannot be fully determined from its
definition; we need to know the context of its use. The dead code -- the
z definition -- although contributing nothing to
f in a `misleading', from the point of view of the live
code, context, thus causing the type error. It is worth contemplating
how come the type error is reported in the `live' expression
rather than in the dead code (even when the dead code appears
after the live one).
It is likely less known that the dead code can have the opposite effect: removing it makes the program untypeable. As far as the computation and its result are concerned, such code does nothing. Still, it does something useful, from the type checker point of view. For example, the following program compiles successfully.
let id = fun u -> u let f = id id let exp x = f x let z () = f TrueIf we remove the definition of
z, it will not: the compiler will blame the definition of
f, saying its type ``contains type variables that cannot be generalized''. Here is a similar example in Haskell98:
main = do x <- return  print (x == x) let z = (x == "") return ()As written, it type-checks and runs. If we delete the definition of the seemingly irrelevant (and also later-bound)
z, the type checker complains, about ``Ambiguous type variable `t0' arising from a use of `==' in x==x''. Sometimes dead may help alive, by constraining them.
Most surprisingly, dead code may affect the result of the program -- even though it is not even executed. Here is a Haskell98 example:
instance Num Char main = do x <- return  print x let y = abs (head x) let z = (x == "") return ()As written, the program prints
"". If we remove the irrelevant
zdefinition, the printed result changes to
. Incidentally, the program has another irrelevant definition, of
y. What happens if that is removed?
Here is an even simpler Haskell98 example, distilling the problem that Andreas Abel's student ran into when doing RGB-conversion in some graphics-related code. It is not just an academic problem.
main = do x <- return (2^65) print (x == 0) let z = x > length  return ()The program prints
True; if we remove the
zdefinition, the result changes to
The context-sensitivity of type inference per se should not be surprising. Type inference, in essence, guesses the type of each subexpression so to make the whole expression typeable. What makes the algorithm deterministic is `lazy guessing': representing the guess by a fresh type variable. As the inference proceeds and more context is seen, something, hopefully, would tell the concrete type to be substituted for the type variable.
The problem behind all our examples is the scope of the introduced type variables. One, `natural' scope is the scope of the current definition: the let-binding. If some type variables are left undetermined at the end of type checking the definition, they are generalized. What if the generalization is impossible/inapplicable? One design choice is to reject the definition as ill-typed. As Andreas Abel indicated, that is what Agda does, in effect. And also MetaOCaml, but for the generated-code variables; as soon as a variable is caught leaving the scope of its intended binder, the code generation process is terminated.
I guess, confining type variables to the scope of program definitions was considered draconian by language designers. They allowed the type variables to float into a larger context, in the hope that the concrete type for them will eventually be found. Therefore, the type of a defined identifier gets determined not only by its definition but also by its use. This is the non-compositionality that some, Andreas Abel in particular, find objectionable. Still, Haskell and OCaml do impose a strict limit on type variables: the compilation unit. If some variables are left unbound and ungeneralized when the whole unit has been type-checked, something has to be done. The OCaml compiler rejects the unit. Haskell attempts to apply the type defaulting rules; if they do not help, the program is declared ambiguous.
Most of the time the non-compositionality is `benign' -- the contexts of use for an identifier, if conflicting, may make the program untypeable, but will not change its result. However, when type defaulting, reflection, overlapping instances, etc. enter the scene, we are up to real surprises.
Refined Environment Classifiers: Type- and Scope-Safe Code Generation with
The global, unrestricted scope of type variables has much in common with variable names in the generated code. The calculus
<NJ> developed in the cited paper uses a special `name heap' to
track such future-stage names.
The code below implements this approach. We use Scheme
notation for the source language (we could just as well supported ML
or Haskell-like notations). The notation for type terms is infix, with
the right-associative arrow. As an example, the end of the file
type-inference.scm shows the derivation for
reset from their types in the
continuation monad. Given the type:
(define (cont a r) `((,a -> . ,r) -> . ,r)) (((a -> . ,(cont 'b 'r)) -> . ,(cont 'b 'b)) -> . ,(cont 'a 'b))within 2 milli-seconds, we obtain the term for shift:
(lambda (_.0) (lambda (_.1) ((_.0 (lambda (_.2) (lambda (_.3) (_.3 (_.1 _.2))))) (lambda (_.4) _.4))))From the Curry-Howard correspondence, determining a term for a type is tantamount to proving a theorem -- in intuitionistic logic as far as our language is concerned. We formulate the proposition in types, for example:
(define (neg x) `(,x -> . F)) (,(neg '(a * b)) -> . ,(neg (neg `(,(neg 'a) + ,(neg 'b)))))This is one direction of the deMorgan law. In intuitionistic logic, deMorgan law is more involved:
NOT (A & B) == NOTNOT (NOT A | NOT B)The system gives us the corresponding term, the proof:
(lambda (_.0) (lambda (_.1) (_.1 (inl (lambda (_.2) (_.1 (inr (lambda (_.3) (_.0 (cons _.2 _.3))))))))))The de-typechecker can also prove theorems in classical logic, via the double-negation (aka CPS) translation. We formulate the proposition:
(neg (neg `(,(neg 'a) + ,(neg (neg 'a)))))and, within 403 ms, obtain its proof:
(lambda (_.0) (_.0 (inr (lambda (_.1) (_.0 (inl _.1))))))The proposition is the statement of the Law of Excluded Middle, in the double-negation translation.
Programming languages may help studying logic, indeed.
let, sums and products -- and its type. The code contains many comments that explain the notation, the incremental composition of the type checking relation and taking the advantage of the first-class status of Kanren relations for type checking of polymorphic
Kanren code that illustrates the application of the inverse typechecker to proving theorems in intuitionistic and classical logics.
Reversing Haskell typechecker
Deriving a program from its incomplete specification such as
type using constraints, background knowledge, heuristics or training
set is the subject of inductive programming
gensyms) a fresh type variable and watches for its scope, following the proof rule for universal introduction. Due to this implementation, rank-2 polymorphism has been used (some would say abused) to express static capabilities and assure a wide variety of safety properties statically, including:
Joint work with Chung-chieh Shan.
Lightweight monadic regions
A few applications of fresh type variables, not counting the ST monad