previous   next   up   top

Approximate Dependent-Type Programming


Haskell is not a dependently-typed language -- at least it was not designed that way. Nevertheless, even Haskell98 admits some form of dependent types, enough to statically prevent the annoying ``head of empty list'' error. After multi-parameter type classes with functional dependencies were introduced, it was quickly (Hallgren, 2001) realized that we can compute with types. HList (2004) has lifted the entire list library to the type level, opening the floodgates of type computation, all the way to SkewLists (Martinez et al., PEPM 2013) and RSA. Later on, type families have made programming with types deceptively more functional. 2004 brought another key realization: polymorphic recursion lets us dynamically select among the family of types and hence makes the selected type in effect depend on a run-time value. Singleton programming and GADTs perfected this pattern. ``This much is clear: many programmers are already finding practical uses for the approximants to dependent types which mainstream functional languages (especially Haskell) admit, by hook or by crook.'' (Altenkirch et al., 2005)

This page is a panorama of approximations to dependently-typed programming. We will see how closely Haskell comes to dependent types, yet not reaching them. No worry: cruder approximations turn out just as, if not more, practically useful as the overweight approaches.

What is a dependent type
Lightweight approaches
The standard approximation (to be written)
Heavy type-level computation


What is a dependent type

An ordinary type such as [a] may depend on other types -- in our case, the types of list elements -- but not on the values of those elements or their number. A dependent type does depend on such dynamic values. For example, in the hypothetical dependent Haskell, the function to produce a list with n copies of a given value could have the following signature
     replicate :: (n::Nat) -> a -> List n a
List n a is the type of lists with elements of type a and the length exactly n, where n is a non-negative integer Nat. To be more precise, List n a is a family of types (of lists of various lengths with elements of type a); the natural number n selects one type of that family: ``indexes within the family''. Such an indexing is one of the common manifestations of dependent types. The index n is the argument of replicate: it is a value that is not known until the run-time, when the function is applied.

The function to append length-indexed lists List n a naturally has the signature:

     append :: List n a -> List m a -> List (n+m) a
where n+m is the ordinary addition of two natural numbers n and m. Thus not only values may appear in types but also arbitrary expressions (terms). The pay-off for including the length of the list in its type is being able to give head and tail the signatures:
     head :: List (succ n) a -> a
     tail :: List (succ n) a -> List n a
that make their applications to the empty list ill-typed. Taking the head/tail of an empty list -- the functional programming equivalent of the infamous NullPointerException -- is hence statically prevented. We no longer have to puzzle out the location of the Prelude.head: empty list exceptions or wonder how to deal with it in the production code. Such an error simply occur in a well-typed program.

Curry-Howard correspondence regards types as propositions and programs as proofs. In simply-typed systems, types (propositions) are made of type constants such as Int (atomic propositions), combined by function arrow (implication), pairing (conjunction), etc. Simply-typed systems hence correspond to propositional logic. First-order predicate logic distinguishes between terms (denoting ``things'': elements of some domain) and propositions, stating properties and relations among terms. Propositions like Prime(n) naturally contain terms, n in our case. Dependent types thus are the Curry-Howard interpretation of the first-order predicate logic. Therefore, any program specification written in first-order logic can be expressed as a dependent type, to be checked by the compiler. ``Dependently typed programs are, by their nature, proof carrying code.'' (Altenkirch et al.)

One should distinguish a dependent type (which depends on a dynamic value) from a polymorphic type such as Maybe a. The type List n a from our running example is indexed by the value (the list length) and by the type (of its elements): it is both dependent (in n) and polymorphic (or, technically, parametric), in a. In Agda, List would have the type Nat -> Set -> Set, which clearly shows the distinction between the dependency and the parametric polymorphism. The first argument of List is a natural number, an element of the domain Nat. The second argument of List is a Set -- itself a proposition. Parametric polymorphism hence has the Curry-Howard correspondence to (a fragment of a) second-order logic. The different nature of the two arguments in List n a has many consequences: since the list length n is a natural number, one may apply to it any operation on natural numbers such as addition, increment, multiplication, etc. -- as we have seen in the append example. On the other hand, there is much less one can do with types. List concatenation and many similar operations keep the type of list elements but change its length.

The main issue of dependently-typed systems is illustrated by the following simple example (which will be our running example):

     pcomm :: Bool -> List n a -> List m a -> List (n+m) a
     pcomm b l1 l2 = if b then append l1 l2 else append l2 l1

The type checker has to verify that append l1 l2 and append l2 l1 in the conditional branches have the same type. Recalling the signature of append, the type checker has to ascertain that n+m is equal to m+n. Now, deciding if two types are the same involves determining if two expressions are equal, which is generally undecidable (think of functions or recursive expressions). There is a bigger problem: in our example, n and m are just variables, whose values will be known only at run-time. The type-checker, which runs at compile-time, therefore has to determine that n+m is equal to m+n without knowing the concrete values of n and m. We know that natural addition is commutative, but the type-checker does not. It is usually not so smart to figure out the commutativity from the definition of addition. Therefore, we have to somehow supply the proof of the commutativity to the type-checker. Programming with dependent types involves a great deal of theorem proving.

Thorsten Altenkirch, Conor McBride, and James McKinna: Why Dependent Types Matter. April 2005.
< >


Non-empty lists

Errors such as taking head or tail of the empty list in Haskell are equivalent to the dereferencing of the zero pointer in C/C++ or NullPointerException in Java. These errors occur because the domain of the function is smaller than the function's type suggests. For example, the type of head says that the function applies to any list. In reality, it can be meaningfully applied only to a non-empty list. One can eliminate such errors by giving functions head and tail a more precise type, such as FullList a. Languages like Cyclone and Cw do exactly that.

We stress that the head-of-empty-list errors can be eliminated now, without any modification to the Haskell type system, without developing any new tool. Already Haskell98 can do that. The same technique applies to OCaml and even Java and C++. The only required advancement is in our thinking and programming style.

Thinking of full lists as a separate type from ordinary, potentially empty lists does affect our programming style -- but it does not have to break the existing code. The new style is easy to introduce gradually. Besides safety, its explicitness makes list processing algorithms more insightful, separating out algorithmically meaningful empty list checks from the redundant safety checks. Let us see some examples.

Assume the following interface

     type FullList a -- assume abstract. For an implementation, see below
     fromFL :: FullList a -> [a]
     indeedFL :: [a] -> w -> (FullList a -> w) -> w  -- an analogue of `maybe'
     headS :: FullList a -> a
     tailS :: FullList a -> [a]
     -- Adding something to a general list surely gives a non-empty list
     infixr 5 !:
     class Listable l where
         (!:) :: a -> l a -> FullList a
All the above functions are total, in particular, headS and tailS. The operation fromFL lets us forget that the list is FullList, giving the ordinary list. The application headS [] obviously does not type-check. Less obvious is that the application headS $ some_expression will also fail to type check unless we can statically assure some_expression produces a FullList. To see how difficult it is to obtain that assurance, let's take an example. First is the accumulating list reversal function, written to explicitly use head and tail:
     regular_reverse :: [a] -> [a]
     regular_reverse l = loop l []
        loop [] accum = accum
        loop l  accum = loop (Prelude.tail l) (Prelude.head l : accum)
Let us re-write it using safe head and tail functions:
     safe_reverse :: [a] -> [a]
     safe_reverse l = loop l [] 
        loop l accum = indeedFL l accum $
                          (\l -> loop (tailS l) (headS l : accum))
     test1 = safe_reverse [1,2,3]
That was straightforward: we relied on indeedFL to perform case analysis on the argument list (if it is empty or not). We had to do this analysis anyway, according to the algorithm. In the case branch for the non-empty list, we statically know that l is non-empty and so the applications of headS and tailS are well-typed. Even the type-checker can see that. Here is another example, which should be self-explanatory.
     safe_append :: [a] -> FullList a -> FullList a
     safe_append [] l    = l
     safe_append (h:t) l = h !: safe_append t l
     l1 :: FullList Int
     l1 = 1 !: 2 !: []
     -- We can apply safe_append on two FullList without any problems
     test5 = tailS $ safe_append (fromFL l1) l1
     -- [2,1,2]

Before further discussing this programming style and the problem of backwards compatibility (``do we have to re-write all the code?''), let us see the implementation. Recall, our goal is to define the type of FullList, of lists that are guaranteed to be non-empty. One can see two approaches: one particular and insightful, and the other is generic and insightful. We start with the first one: a non-empty list for sure has at least one element, to be carried around explicitly:

     data FullList a = FullList a [a]  -- carry the list head explicitly
     fromFL :: FullList a -> [a]
     fromFL (FullList x l) = x : l
     headS :: FullList a -> a
     headS (FullList x _) = x
     tailS :: FullList a -> [a]
     tailS (FullList _ x) = x
The other functions of our interface are equally straightforward, see the source code for details. The undisputed advantage of the implementation is its obvious correctness: FullList represents a non-empty list by its very construction, without any pre-conditions and reservations. FullList truly is a non-empty list. There is a deep satisfaction in finding a data structure that ensures a desired property by its very construction, where the property is built-in. In this respect, one can't help but think of Chris Okasaki's work.

There is another approach of representing FullList, which easily generalizes to other structures.

     module NList (FullList, fromFL, headS, tailS, ...) where
     newtype FullList a = FullList [a]  -- data constructor is not exported!
     fromFL (FullList x) = x
     -- The following are _total_ functions
     -- They are guaranteed to be safe, and so we could have used
     -- unsafeHead# and unsafeTail# if GHC provided them.
     headS :: FullList a -> a
     headS (FullList (x:_)) = x
     tailS :: FullList a -> [a]
     tailS (FullList (_:x)) = x
We introduce the abstract data type FullList -- abstract in the sense that its constructor, also named FullList, is not exported. The only way to construct and manipulate the values of that type is to use the operations exported by the module. All the exported operations ensure that FullList represents a non-empty list.

One may regard the abstract type FullList as standing for the proposition (invariant) that represented list is non-empty. Now we have something to prove: we have to verify, manually or semi-automatically, that all operations within the module NList whose return type is FullList respect the invariant and ensure the truth of the non-emptiness proposition. Once we have verified these exported constructors, all operations that consume FullList, within NList or outside, can take this non-emptiness proposition for granted. Therefore, we are justified in using unsafe-head operations in implementing headS. Compared to the first implementation, the fact that FullList represents a non-empty list is no longer obvious and has to be proven. Fortunately, we only have to prove the operations within NList, that is, the ones that make use of the data constructor FullList. All other functions, which produce FullList merely by invoking the operations of NList, ensure the non-emptiness invariant by construction and do not need a proof. The advantage of the second implementation is the easy generalization to ByteString-like packed lists, etc. The data construction FullList is merely a newtype wrapper, with no run-time overhead. Thus the second implementation provides the safety of head and tail operations without sacrificing efficiency.

In the old (2006) discussion of non-empty lists on Haskell-Cafe, Jan-Willem Maessen wrote: ``In addition, we have this rather nice assembly of functions which work on ordinary lists. Sadly, rewriting them all to also work on NonEmptyList or MySpecialInvariantList is a nontrivial task.'' Backwards compatibility is indeed a serious concern: no matter how better a new programming style may be, the mere thought of re-writing the existing code is a deterrent. Suppose we have a function foo:: [a] -> [a] (whose code, if available, we'd rather not change) and we want to write something like

     	\l -> [head l, head (foo l)]
The first attempt of using the safe functions
     	\l -> indeedFL l onempty (\l -> [headS l, headS (foo l)])
does not type: foo applies to [a] rather than FullList a; furthermore, the result of foo is not FullList a, required by headS. The first problem is easy to solve: FullList a can always be cast into the general list, with fromFL. We insist on writing the latter function explicitly, which keeps the type system simple, free of subtyping and implicit coercions. One may regard fromFL as an analogue of fromIntegral -- which, too, we have to write explicitly, in any code with more than one sort of integral numbers (e.g., Int and Integer, or Int and CInt).

If we are not sure if our foo maps non-empty lists to non-empty lists, we really should handle the empty list case:

     	\l -> indeedFL l onempty $
     	       \l -> [headS l, indeedFL (foo $ fromFL l) onempty' headS]
If we have a hunch that foo indeed maps non-empty lists to non-empty lists, but we are too busy to verify it, we can write
     \l -> indeedFL l onempty $
            \l -> [headS l, indeedFL (foo $ fromFL l) (error msg) headS]
       where msg = "I'm quite sure foo maps non-empty lists to " ++
                   "non-empty lists. I'll be darned if it doesn't."
That would get the code running. Possibly at some future date (during the code review?) we will be called to justify the hunch, to whatever required degree of formality (informal argument, formal proof). If we fail at this justification, we'd better think what to do if the result of foo turns out empty. If we succeed, we would be given permission to add to the module NList the following definition:
     	nfoo (FullList x) = FullList $ foo x
after which we can write
     	\l -> indeedFL l onempty (\l -> [headS l, headS (nfoo l)])
with no extra empty list checks.

In conclusion, we have demonstrated the programming style that ensures safety without sacrificing efficiency. The key idea is that an abstract data type ensures (possibly quite sophisticated) propositions about the data -- so long as the very limited set of basic constructors satisfy the propositions. This main idea is very old, advocated by Milner and Morris in the mid-1970s. If there is a surprise in this, it is in the triviality of approach. One can't help but wonder why we do not program in this style.

The current version is August 2015; Original: November 2006.
NList0.hs [2K]
NList.hs [2K]
NListTest.hs [2K]
The two complete implementations of the library and the tests

James H. Morris Jr.: Protection in Programming Languages Comm. of the ACM, 1973, V16, N1, pp. 15-21
< >

Lightweight dependent typing
Longer explanations of the technique, justification, formalization, and more complex examples

The FullList library was first presented and discussed on Haskell-Cafe in November 2006 and later summarized on Haskell Wiki
< >
The present article is the expanded and elaborated version.


Implicit configurations -- or, type classes reflect the values of types

The configurations problem is to propagate run-time preferences throughout a program, allowing multiple concurrent configuration sets to coexist safely under statically guaranteed separation. This problem is common in all software systems, but particularly acute in Haskell, where currently the most popular solution relies on unsafe operations and compiler pragmas.

We solve the configurations problem in Haskell using only stable and widely implemented language features like the type-class system. In our approach, a term expression can refer to run-time configuration parameters as if they were compile-time constants in global scope. Besides supporting such intuitive term notation and statically guaranteeing separation, our solution also helps improve the program's performance by transparently dispatching to specialized code at run-time. We can propagate any type of configuration data -- numbers, strings, IO actions, polymorphic functions, closures, and abstract data types. No previous approach to propagating configurations implicitly in any language provides the same static separation guarantees.

The enabling technique behind our solution is to propagate values via types, with the help of polymorphic recursion and higher-rank polymorphism. The technique essentially emulates local type-class instance declarations while preserving coherence. Configuration parameters are propagated throughout the code implicitly as part of type inference rather than explicitly by the programmer. Our technique can be regarded as a portable, coherent, and intuitive alternative to implicit parameters. It motivates adding local instances to Haskell, with a restriction that salvages principal types.

Joint work with Chung-chieh Shan.

The current version is August 2004.
tr-15-04.pdf [252K]
Technical report TR-15-04, Division of Engineering and Applied Sciences, Harvard University
This is the expanded version of the paper, with extra appendices.
The original paper was published in the Proceedings of the ACM SIGPLAN 2004 workshop on Haskell
Snowbird, Utah, USA -- September 22, 2004 -- ACM Press, pp. 33 - 44.
Typo: several occurrences of term4' on page 9 should be read as test4'.

Prepose.hs [10K]
The complete code for the paper, updated for the current GHC

Last updated August 11, 2015

This site's top page is
Your comments, problem reports, questions are very welcome!

Converted from HSXML by HSXML->HTML