Lens as a Divisibility Relation: Goofin’ Off With the Algebra of Types

Types have an algebra very analogous to the algebra of ordinary numbers (video). This is the basic table of correspondences. Code with all the extensions available here.

One way to see that this makes sense is by counting the cardinality of types built out of these combinators. Unit is the type with 1 inhabitant. Void has 0 inhabitants. If a has n and b has m possible values, then Either a b has n + m inhabitants, (a,b) has n*m and there are n^m possible tabulations of a->b. We’re gonna stick to just polynomials for the rest of this, ignoring a->b.

Another way of looking at this is if two finitely inhabited types have the same number of inhabitants, then the types can be put into an isomorphism with each other. In other words, types modulo isomorphisms can be thought as representing the natural numbers. Because of this, we can build a curious proof system for the natural numbers using ordinary type manipulation.

In addition, we also get a natural way of expressing and manipulating polynomials.Polymorphic types can be seen as being very similar to polynomial expressions with natural coefficients N[x]. The polymorphic type variables have the ability to be instantiated to any value, corresponding to evaluating the polynomial for some number.

The Lens ecosystem gives some interesting combinators for manipulating this algebra. The type Iso' a b contains isomorphisms. Since we’re only considering types up to isomorphism, this Iso' represents equality. We can give identity isomorphisms, compose isomorphisms and reverse isomorphisms.

We can already form a very simple proof.

Now we’ll add some more combinators, basically the axioms that the types mod isos are a commutative semiring. Semirings have an addition and multiplication operator that distribute over each other. It is interesting to note that I believe all of these Iso' actually are guaranteed to be isomorphisms ( to . from = id and from . to = id ) because of parametricity. from and to are unique ignoring any issues with bottoms because the polymorphic type signature is so constraining. This is not usually guaranteed to be true in Haskell just from saying it is an Iso'. If I give you an Iso' Bool Bool it might actually be the iso (const True) (const True) for example, which is not an isomorphism.

There are also combinators for lifting isomorphisms into bifunctors: firsting, seconding, and bimapping. These are important for indexing into subexpressions of our types in a point-free style.

Here is a slightly more complicated proof now available to us.

We can attempt a more interesting and difficult proof. I developed this iteratively using . _ hole expressions so that GHC would tell me what I had manipulated my type to at that point in my proof.

Artwork Courtesy of David. Sorry for any motion sickness.

The proof here is actually pretty trivial and can be completely automated away. We’ll get to that later.

If Iso' is equality, what about the other members of the Lens family? Justin Le says that Lens s a are witness to the isomorphism of a type s to the tuple of something and a. Prism witness a similar thing for sums. Once we are only considering types mod isos, if you think about it, these are expressions of two familiar relations on the natural numbers: the inequality relation and the divisibility relation

Mathematically, these relations can be composed with equalities, just like in the lens hierarchy Lens and Prism can be composed with Iso. Both form a category, since they both have id and (.).

Here are a couple identities that we can’t derive from these basic combinators. There are probably others. Woah-ho, my bad. These are totally derivable using id_mul, id_plus, mul_zero, _1, _2, _Left, _Right.

Pretty neat! Random thoughts and questions before we get into the slog of automation:

  • Traversal is the “is polynomial in” relation, which seems a rather weak statement on it’s own.
  • Implementing automatic polynomial division is totally possible and interesting
  • What is the deal with infinite types like [a]? Fix. I suppose this is a theory of formal power series / combinatorial species. Fun combinatorics, generatingfunctionology. Brent Yorgey did his dissertation on this stuff. Wow. I’ve never really read this. It is way more relevant than I realized.
  • Multivariate polynomial algorithms would also be interesting to explore (Grobner basis, multivariate division)
  • Derivatives of types and zippers – Conor McBride
  • Negative Numbers and Fractions?
  • Lifting to rank-1 types. Define Negative and Fractions via Galois connection?

Edit: /u/carette (wonder who that is šŸ˜‰ ) writes:

“You should dig into
J Carette, A Sabry Computing with semirings and weak rig groupoids, in Proceedings of ESOP 2016, p. 123-148. Agda code in https://github.com/JacquesCarette/pi-dual/tree/master/Univalence. A lot of the algebra you develop is there too.

If you hunt around in my repos, you’ll also find things about lenses, exploring some of the same things you mention here.”

Similar ideas taken further and with more sophistication. Very interesting. Check it out.


Our factor example above was quite painful, yet the theorem was exceedingly obvious by expansion of the left and right sides. Can we automate that? Yes we can!

Here’s the battle plan:

  • Distribute out all expressions like a*(b+c) so that all multiplication nodes appear at the bottom of the tree.
  • Reduce the expression by absorbing all stupid a*1, a*0, a+0 terms.
  • Reassociate everything to the right, giving a list like format
  • Sort the multiplicative terms by power of the variable

Once we have these operations, we’ll combine them into a canonical operation. From there, most equality proofs can be performed using the rewrite operation, which merely puts both sides into canonical form

Once we have those, the painful theorem above and harder ones becomes trivial.

Now we’ll build the Typeclasses necessary to achieve each of these aims in order. The Typeclass system is perfect for what we want to do, as it builds terms by inspecting types. It isn’t perfect in the sense that typeclass pattern matching needs to be tricked into doing what we need. I have traded in cleverness and elegance with verbosity.

In order to make our lives easier, we’ll need to tag every variable name with a newtype wrapper. Otherwise we won’t know when we’ve hit a leaf node that is a variable. I’ve used this trick before here in an early version of my faking Compiling to Categories series. These wrappers are easily automatically stripped.

A common pattern I exploit is to use a type family to drive complicated recursion. Closed type families allow more overlap and default patterns which is very useful for programming. However, type families do not carry values, so we need to flip flop between the typeclass and type family system to achieve our ends.

Here is the implementation of the distributor Dist. We make RDist and LDist typeclasses that make a sweep of the entire tree, using ldist and rdist as makes sense. It was convenient to separate these into two classes for my mental sanity. I am not convinced even now that I have every case. Then the master control class Dist runs these classes until any node that has a (*) in it has no nodes with (+) underneath, as checked by the HasPlus type family.

Next is the Absorb type class. It is arranged somewhat similarly to the above. Greedily absorb, and keep doing it until no absorptions are left. I think that works.

The Associators are a little simpler. You basically just look for the wrong association pattern and call plus_assoc or mul_assoc until they don’t occur anymore, then recurse. We can be assured we’re always making progress if we either switch some association structure or recurse into subparts.

Finally, the SortTerm routine. SortTerm is a bubble sort. The typeclass Bubble does a single sweep of swapping down the type level list-like structure we’ve built. The SortTerm uses the Sorted type family to check if it is finished. If it isn’t, it call Bubble again.

Hope you thought this was neat!

Proving Addition is Commutative in Haskell using Singletons

Yesterday morning, I was struck with awe at how amazingly cool the dependently typed approach (Agda, Idris, Coq) to proving programs is. It is great that I can still feel that after tinkering around with it for a couple years. It could feel old hat.

In celebration of that emotion, I decided to write a little introductory blog post about how to do some proving like that in Haskell. Haskell can replicate this ability to prove with varying amounts of pain. For the record, there isn’t anything novel in what follows.

One methodology for replicating the feel of dependent types is to use the singletons pattern. The singleton pattern is a way of building a data type so that there is an exact copy of the value of a variable in its type.

For future reference, I have on the following extensions, some of which I’ll mention when they come up.

Here’s how the pattern goes.

Step 1: define your ordinary data type. Bool is already defined in the Prelude, but here is what it looks like

Step 2: turn on the DataKinds extension. This automatically promotes any data type constructor like True or False or Just into types that have apostrophes in front of them 'True, 'False, 'Just. This is mostly a convenience. You could have manually defined something similar like so

Step 3: Define your singleton datatype. The singleton datatype is a GADT (generalized abstract data type). GADT declarations take on a new syntax. It helps to realize that ordinary type constructors like Just are just functions. You can use them anywhere you can use functions. Just has the type a -> Maybe a. It might help to show that you can define a lower case synonym.

just is clearly just a regular function. What makes constructors a special thing (not quite “just functions”) is that you can also pattern match on them. Data constructors are functions that “do nothing”. They hold the data until eventually you get it back again via pattern matching.

So why don’t you write the type signature of the constructors when you define them? Why does using a data statement look so different than a function definition? The GADT syntax brings the two concepts visually closer.

Letting you define the type signature for the constructor let’s you define a constrained type signature, rather than the inferred most general one. It’s similar to the following situation. If you define an identity function id, it has the polymorphic type a -> a. However, you can explicitly constrain the function with an identical implementation. If you try to use boolid on an Int it is a type error.

The GADT syntax let’s you constrain what the type signature of the constructor is, but also very interestingly, let’s the type checker infer information when you pattern match into the GADT.

With all that spiel, here is the singleton type for Bool as a GADT.

You have made an exact copy of the value at the type level. When you pattern match into a variable x of type SBool s in the STrue branch, it knows that s ~ 'True and in the SFalse branch it knows that s ~ 'False.

Here’s the analogous construction for a Peano natural number

Now let’s talk about programming.

Addition is straightforward to define for our Nat.

Let’s replicate this definition at the type level. The extension we’ll want is TypeFamilies. Type families enables a syntax and feature for defining functions on types very similarly to how you define regular functions.

Now finally, we can exactly mirror this definition on singletons

In the type signature SNat is kind of like a weirdo forall. It is a binding form that generates a new type variable you need to express the typelevel connections you want. The type variable n is a typelevel thing that represents the value.

Now let’s talk about proving. Basically, if you’re intent is proving things, I think it is simplest if you forget that the original data type ever existed. It is just a vestigial appendix that makes the DataKinds you need. Only work with singletons. You will then need to make a safe layer translating into and out of the singletons if you want to interface with non-singleton code.

We’re going to want to prove something about equality. The standard definition of equality is

I put the 1 there just so I wouldn’t clash with the Eq typeclass. It’s ugly, sorry. You can find an identical definition in base http://hackage.haskell.org/package/base- that uses the extension TypeOperators for a much cleaner syntax.

Why is this a reasonable equality? You can construct using Refl only when you are just showing that a equals itself. When you pattern match on Refl, the type checker is learning that a ~ b. It’s confusing. Let’s just try using it.

We can prove a couple facts about equality. First off that it is a symmetric relation. If a = b then b = a.

When we pattern match and expose the incoming Refl, the type checker learns that a ~ b in this branch of the pattern match. Now we need to return an Eq1 b a. But we know that a ~ b so this is the same as an Eq1 a a. Well, we can easily do that with a Refl.

Similarly we can prove the transitivity of equality.

Pattern matching on the first equality tells the type checker that a ~ b, the second that b ~ c. Now we need to return a Eq1 a c but this is the same as Eq1 a a because of the ~ we have, so Refl suffices.

It’s all damn clever. I wouldn’t have come up with it.

Now let’s start proving things about our addition operator. Can we prove that

This one is straightforward since obviously 'Zero is 'Zero. How about something slightly more complicated 1 + 0 = 1.

The typechecker can evaluate addition on concrete numbers to confirm this all works out.

Here’s a much more interesting property \forall x. 0 + x = x

This one is also straightforward to prove. Looking at the definition of NPlus knowing that the first argument is 'Zero is enough to evaluate forward.

Here’s our first toughy. \forall x. x + 0 = x This may seem silly, but our definition of NPlus did not treat the two arguments symmetrically. it only pattern match on the first. Until we know more about x, we can’t continue. So how do we learn more? By pattern matching and looking at the possible cases of x.

The first case is very concrete and easy to prove. The second case is more subtle. We learn that x ~ 'Succ x1 for some x1 when we exposed the SSucc constructor. Hence we now need to prove Eq1 (NPlus ('Succ x1) 'Zero) ('Succ x1). The system now has enough information to evaluate NPlus a bit, so actually we need Eq1 ('Succ (NPlus x1 'Zero)) ('Succ x1). The term (NPlus x1 'Zero) looks very similar to what we were trying to prove in the first case. We can use a recursive call to get an equality proof that we pattern match to a Refl to learn that(NPlus x1 'Zero) ~ x1 which will then make the required result Eq1 ('Succ x1) ('Succ x1) which is obviously true via Refl. I learned a neat-o syntax for this second pattern match, called pattern guards. Another way of writing the same thing is

Follow all that? Haskell is not as friendly a place to do this as Idris or Agda is.

Now finally, the piece de resistance, the commutativity of addition, which works in a similar but more complicated way.

A question: to what degree does this prove that nplus or snplus are commutative? The linkage is not perfect. snplus is type constrained to return the same result as NPlus which feels ok. nplus is syntactically identical to the implementation of the other two, but that is the only link, there is nothing compiler enforced.

The existence of non-termination in Haskell also makes everything done here much less fool-proof. It wouldn’t be all that hard to accidentally make a recursive call in one of our proofs that is non terminating and the compiler would accept it and say nothing.

I recommend you check out the links below for more.

Source available here https://github.com/philzook58/singleberg






A Little Bloop on Typed Template Haskell

I’ve found looking at my statistics that short, to the point, no bull crap posts are the most read and probably most useful.

I’ve been tinkering around with typed template Haskell, which a cursory internet search doesn’t make obvious even exists. The first thing to come up is a GHC implementation wiki that seems like it might be stalled in some development branch. No. Typed template Haskell is already in GHC since GHC 7.8. And the first thing that comes up on a template haskell google search is the style of Template Haskell where you get deep into the guts of the syntax tree, which is much less fun.

Here’s a basic addition interpreter example.

The typed splice $$ takes a out of TExpQ a. The typed quote [|| ||] puts it in. I find that you tend to be able to follow the types to figure out what goes where. If you’re returning a TExpQ, you probably need to start a quote. If you need to recurse, you often need to splice. Etc.

You tend to have to put the actual use of the splice in a different file. GHC will complain otherwise.

At the top of your file put this to have the template haskell splices dumped into a file

Or have your package.yaml look something like this

If you’re using stack, you need to dive into .stack/dist/x86/Cabal/build and then /src or into the executable folder something-exe-tmp/app to find .dump-splices files.

Nice. I don’t know whether GHC might have optimized this thing down anyhow or not, but we have helped the process along.

Some more examples: unrolling the recursion on a power function (a classic)

You can unroll a fibonacci calculation

This is blowing up in calculation though (no memoization, you silly head). We can implement sharing in the code using let expression (adapted from https://wiki.haskell.org/The_Fibonacci_sequence). Something to think about.

Tinkering around, you’ll occasionally find GHC can’t splice and quote certain things. This is related to cross stage persistence and lifting, which are confusing to me. You should look in the links below for more info. I hope I’ll eventually get a feel for it.

If you want to see more of my fiddling in context to figure out how to get stuff to compile here’s my github that I’m playing around in https://github.com/philzook58/staged-fun

Useful Link Dump:

Simon Peyton Jones has a very useful talk slides. Extremely useful https://www.cl.cam.ac.uk/events/metaprog2016/Template-Haskell-Aug16.pptx

Matthew Pickering’s post keyed me into that Typed Template Haskell is even there.


MuniHac 2018: Keynote: Beautiful Template Haskell

https://markkarpov.com/tutorial/th.html Mark Karpov has a useful Template Haskell tutorial

Oleg Kiselyov: I’m still trying to unpack the many things that are going on here. Oleg’s site and papers are a very rich resource.

I don’t think that staged metaprogramming requires finally tagless style, as I first thought upon reading some of his articles. It is just very useful.



Typed Template Haskell is still unsound (at least with the full power of the Q templating monad) http://okmij.org/ftp/meta-programming/#TExp-problem

You can write things in a finally tagless style such that you can write it once and have it interpreted as staged or in other ways. There is a way to also have a subclass of effects (Applicatives basically?) less powerful than Q that is sound.

A Touch of Topological Quantum Computation 3: Categorical Interlude

Welcome back, friend.

In the last two posts, I described the basics of how to build and manipulate the Fibonacci anyon vector space in Haskell.

As a personal anecdote, trying to understand the category theory behind the theory of anyons is one of the reasons I started learning Haskell. These spaces are typically described using the terminology of category theory. I found it very frustrating that anyons were described in an abstract and confusing terminology. I really wondered if people were just making things harder than they have to be. I think Haskell is a perfect playground to clarify these constructions. While the category theory stuff isn’t strictly necessary, it is interesting and useful once you get past the frustration.

Unfortunately, I can’t claim that this article is going to be enough to take you from zero to categorical godhood

but I hope everyone can get something out of it. Give it a shot if you’re interested, and don’t sweat the details.

The Aroma of Categories

I think Steve Awodey gives an excellent nutshell of category theory in the introductory section to his book:

“What is category theory? As a first approximation, one could say that category theory is the mathematical study of (abstract) algebras of functions. Just as group theory is the abstraction of the idea of a system of permutations of a set or symmetries of a geometric object, so category theory arises from the idea of a system of functions among some objects.”

For my intuition, a category is any “things” that plug together. The “in” of a thing has to match the “out” of another thing in order to hook them together. In other words, the requirement for something to be a category is having a notion of composition. The things you plug together are called the morphisms of the category and the matching ports are the objects of the category. The additional requirement of always having an identity morphism (a do-nothing connection wire) is usually there once you have composition, although it is good to take especial note of it.

Category theory is an elegant framework for how to think about these composing things in a mathematical way. In my experience, thinking in these terms leads to good abstractions, and useful analogies between disparate things.

It is helpful for any abstract concept to list some examples to expose the threads that connect them. Category theory in particular has a ton of examples connecting to many other fields because it is a science of analogy. These are the examples of categories I usually reach for. Which one feels the most comfortable to you will depend on your background.

  • Hask. Objects are types. Morphisms are functions between those types
  • Vect. Objects are vector spaces, morphisms are linear maps (roughly matrices).
  • Preorders. Objects are values. Morphisms are the inequalities between those values.
  • Sets. Objects are Sets. Morphisms are functions between sets.
  • Cat. Objects are categories, Morphisms are functors. This is a pretty cool one, although complete categorical narcissism.
  • Systems and Processes.
  • The Free Category of a directed graphs. Objects are vertices. Morphisms are paths between vertices

Generic Programming and Typeclasses

The goal of generic programming is to run programs that you write once in many way.

There are many ways to approach this generic programming goal, but one way this is achieved in Haskell is by using Typeclasses. Typeclasses allow you to overload names, so that they mean different things based upon the types involved. Adding a vector is different than adding a float or int, but there are programs that can be written that reasonably apply in both situations.

Writing your program in a way that it applies to disparate objects requires abstract ways of talking about things. Mathematics is an excellent place to mine for good abstractions. In particular, the category theory abstraction has demonstrated itself to be a very useful unified vocabulary for mathematical topics. I, and others, find it also to be a beautiful aesthetic by which to structure programs.

In the Haskell base library there is a Category typeclass defined in base. In order to use this, you need to import the Prelude in an unusual way.

The Category typeclass is defined on the type that corresponds to the morphisms of the category. This type has a slot for the input type and a slot for the output type. In order for something to be a category, it has to have an identity morphisms and a notion of composition.

The most obvious example of this Category typeclass is the instance for the ordinary Haskell function (->). The identity corresponds to the standard Haskell identity function, and composition to ordinary Haskell function composition.

Another example of a category that we’ve already encountered is that of linear operators which we’ll call LinOp. LinOp is an example of a Kliesli arrow, a category built using monadic composition rather than regular function composition. In this case, the monad Q from my first post takes care of the linear pipework that happens between every application of a LinOp. The fish <=< operator is monadic composition from Control.Monad.

A related category is the FibOp category. This is the category of operations on Fibonacci anyons, which are also linear operations. It is LinOp specialized to the Fibonacci anyon space. All the operations we've previously discussed (F-moves, braiding) are in this category.

The "feel" of category theory takes focus away from the objects and tries to place focus on the morphisms. There is a style of functional programming called "point-free" where you avoid ever giving variables explicit names and instead use pipe-work combinators like (.), fst, snd, or (***). This also has a feel of de-emphasizing objects. Many of the combinators that get used in this style have categorical analogs. In order to generically use categorical typeclasses, you have to write your program in this point free style.

It is possible for a program written in the categorical style to be a reinterpreted as a program, a linear algebra operation, a circuit, or a diagram, all without changing the actual text of the program. For more on this, I highly recommend Conal Elliot's  compiling to categories, which also puts forth a methodology to avoid the somewhat unpleasant point-free style using a compiler plug-in. This might be an interesting place to mine for a good quantum programming language. YMMV.

Monoidal Categories.

Putting two processes in parallel can be considered a kind of product. A category is monoidal if it has this product of this flavor, and has isomorphisms for reassociating objects and producing or consuming a unit object. This will make more sense when you see the examples.

We can sketch out this monoidal category concept as a typeclass, where we use () as the unit object.


In Haskell, the standard monoidal product for regular Haskell functions is (***) from Control.Arrow. It takes two functions and turns it into a function that does the same stuff, but on a tuple of the original inputs. The associators and unitors are fairly straightforward. We can freely dump unit () and get it back because there is only one possible value for it.

The monoidal product we'll choose for LinOp is the tensor/outer/Kronecker product.

Otherwise, LinOp is basically a monadically lifted version of (->). The one dimensional vector space Q () is completely isomorphic to just a number. Taking the Kronecker product with it is basically the same thing as scalar multiplying (up to some shuffling).

Now for a confession. I made a misstep in my first post. In order to make our Fibonacci anyons jive nicely with our current definitions, I should have defined our identity particle using type Id = () rather than data Id. We'll do that now. In addition, we need some new primitive operations for absorbing and emitting identity particles that did not feel relevant at that time.

With these in place, we can define a monoidal instance for FibOp. The extremely important and intriguing F-move operations are the assoc operators for the category. While other categories have assoc that feel nearly trivial, these F-moves don't feel so trivial.

This is actually useful

The parC operation is extremely useful to explicitly note in a program. It is an opportunity for optimization. It is possible to inefficiently implement parC in terms of other primitives, but it is very worthwhile to implement it in new primitives (although I haven't here). In the case of (->), parC is an explicit location where actual computational parallelism is available. Once you perform parC, it is not longer obviously apparent that the left and right side of the tuple share no data during the computation. In the case of LinOp and FibOp, parC is a location where you can perform factored linear computations. The matrix vector product (A \otimes B)(v \otimes w) can be performed individually (Av)\otimes (Bw). In the first case, where we densify A \otimes B and then perform the multiplication, it costs O((N_A N_B)^2) time, whereas performing them individually on the factors costs O( N_A^2 + N_B^2) time, a significant savings. Applied category theory indeed.


Judge Dredd courtesy of David

Like many typeclasses, these monoidal morphisms are assumed to follow certain laws. Here is a sketch (for a more thorough discussion check out the wikipedia page):

  • Functions with a tick at the end like assoc' should be the inverses of the functions without the tick like assoc, e.g. assoc . assoc' = id
  • The parC operation is (bi)functorial, meaning it obeys the commutation law parC (f . f') (g . g') = (parC f g) . (parC f' g') i.e. it doesn't matter if we perform composition before or after the parC.
  • The pentagon law for assoc: Applying leftbottom is the same as applying topright
  • The triangle law for the unitors:

String Diagrams

String diagrams are a diagrammatic notation for monoidal categories. Morphisms are represented by boxes with lines.

Composition g . f is made by connecting lines.

The identity id is a raw arrow.

The monoidal product of morphisms f \otimes g is represented by placing lines next to each other.

The diagrammatic notion is so powerful because the laws of monoidal categories are built so deeply into it they can go unnoticed. Identities can be put in or taken away. Association doesn't even appear in the diagram. The boxes in the notation can naturally be pushed around and commuted past each other.

This corresponds to the property

(id \otimes g) \circ (f \otimes id) = (f \otimes id) \circ (id \otimes g)

What expression does the following diagram represent?

Is it (f \circ f') \otimes (g \circ g') (in Haskell notation parC (f . f') (g . g') )?

Or is it (f \otimes g) \circ (f' \otimes g') (in Haskell notation (parC f g) . (parC f' g')?

Answer: It doesn't matter because the functorial requirement of parC means the two expressions are identical.

There are a number of notations you might meet in the world that can be interpreted as String diagrams. Three that seem particular pertinent are:

  • Quantum circuits
Image result for quantum circuits
  • Anyon Diagrams!

Braided and Symmetric Monoidal Categories: Categories That Braid and Swap

Some monoidal categories have a notion of being able to braid morphisms. If so, it is called a braided monoidal category (go figure).

The over and under morphisms are inverse of each other over . under = id. The over morphism pulls the left morphism over the right, whereas the under pulls the left under the right. The diagram definitely helps to understand this definition.

These over and under morphisms need to play nice with the associator of the monoidal category. These are laws that valid instance of the typeclass should follow. We actually already met them in the very first post.

If the over and under of the braiding are the same the category is a symmetric monoidal category. This typeclass needs no extra functions, but it is now intended that the law over . over = id is obeyed.

When we draw a braid in a symmetric monoidal category, we don't have to be careful with which one is over and under, because they are the same thing.

The examples that come soonest to mind have this symmetric property, for example (->) is a symmetric monoidal category..

Similarly LinOp has an notion of swapping that is just a lifting of swap

However, FibOp is not symmetric! This is perhaps at the core of what makes FibOp so interesting.

Automating Association

Last time, we spent a lot of time doing weird typelevel programming to automate the pain of manual association moves. We can do something quite similar to make the categorical reassociation less painful, and more like the carefree ideal of the string diagram if we replace composition (.) with a slightly different operator

Before defining reassoc, let's define a helper LeftCollect typeclass. Given a typelevel integer n, it will reassociate the tree using a binary search procedure to make sure the left branch l at the root has Count l = n.

Once we have LeftCollect, the typeclass ReAssoc is relatively simple to define. Given a pattern tree, we can count the elements in it's left branch and LeftCollect the source tree to match that number. Then we recursively apply reassoc in the left and right branch of the tree. This means that every node has the same number of children in the tree, hence the trees will end up in an identical shape (modulo me mucking something up).

It seems likely that one could write equivalent instances that would work for an arbitrary monoidal category with a bit more work. We are aided somewhat by the fact that FibOp has a finite universe of possible leaf types to work with.

Closing Thoughts

While our categorical typeclasses are helpful and nice, I should point out that they are not going to cover all the things that can be described as categories, even in Haskell. Just like the Functor typeclass does not describe all the conceptual functors you might meet. One beautiful monoidal category is that of Haskell Functors under the monoidal product of Functor Composition. More on this to come, I think. https://parametricity.com/posts/2015-07-18-braids.html

We never even touched the dot product in this post. This corresponds to another doodle in a string diagram, and another power to add to your category. It is somewhat trickier to work with cleanly in familiar Haskell terms, I think because (->) is at least not super obviously a dagger category?

You can find a hopefully compiling version of all my snippets and more in my chaotic mutating Github repo https://github.com/philzook58/fib-anyon

See you next time.


The Rosetta Stone paper by Baez and Stay is probably the conceptual daddy of this entire post (and more).

Bartosz Milewski's Category Theory for Programmer's blog (online book really) and youtube series are where I learned most of what I know about category theory. I highly recommend them (huge Bartosz fanboy).

Catsters - https://byorgey.wordpress.com/catsters-guide-2/



There are fancier embeddings of category theory and monoidal categories than I've shown here. Often you want constrained categories and the ability to choose unit objects. I took a rather simplistic approach here.




Applicative Bidirectional Programming and Automatic Differentiation

I got referred to an interesting paper by a comment of /u/syrak.


Applicative bidirectional programming (PDF), by Kazutaka Matsuda and Meng Wang

In it, they use a couple interesting tricks to make Lens programming more palatable. Lens often need to be be programmed in a point free style, which is rough, but by using some combinators, they are able to program lenses in a pointful style (with some pain points still left over). It is a really interesting, well-written paper. Lots ‘o’ Yoneda and laws. I’m not doing it justice. Check it out!

A while back I noted that reverse mode auto-differentiation has a type very similar to a lens and in fact you can build a working reverse mode automatic differentiation DSL out of lenses and lens-combinators. Many things about lenses, but not all, transfer over to automatic differentiation. The techniques of Matsuda and Wang do appear to transfer fairly well.

This is interesting to me for another reason. Their lift2 and unlift2 functions remind me very much of my recent approach to compiling to categories. The lift2 function is fanning a lens pair. This is basically what my FanOutput typeclass automated. unlift2 is building the input for a function function by supplying a tuple of projection lenses. This is what my BuildInput typeclass did. I think their style may extend many monoidal cartesian categories, not just lenses.

One can use the function b -> a in many of the situations one can use a in. You can do elegant things by making a Num typeclass of b -> a for example. This little fact seems to extend to other categories as well. By making a Num typeclass for Lens s a when a is a Num, we can use reasonable looking notation for arithmetic.

They spend some time discussing the necessity of a Poset typeclass. For actual lawful lenses, the dup implementation needs a way to recombine multiple adjustments to the same object. In the AD-lens case, dup takes care of this by adding together the differentials. This means that everywhere they needed an Eq typeclass, we can use a Num typeclass. There may be usefulness to building a wrapped type data NZ a = Zero | NonZero a like their Tag type to accelerate the many 0 values that may be propagating through the system.

Unfortunately, as is, the performance of this is abysmal. Maybe there is a way to fix it? Unlifting and lifting destroys a lot of sharing and often necessitates adding many redundant zeros. Why are you doing reverse mode differentiation unless you care about performance? Forward mode is simpler to implement. In the intended use case of Matsuda and Wang, they are working with actual lawful lenses, which have far less computational content than AD-lenses. Good lawful lenses should just be shuffling stuff around a little. Maybe one can hope GHC is insanely intelligent and can optimize these zeros away. One point in favor of that is that our differentiation is completely pure (no mutation). Nevertheless, I suspect it will not without help. Being careful and unlifting and lifting manually may also help. In principle, I think the Lensy approach could be pretty fast (since all it is is just packing together exactly what you need to differentiate into a data type), but how to make it fast while still being easily programmable? It is also nice that it is pretty simple to implement. It is the simplest method that I know of if you needed to port operable reverse mode differentiation to a new library (Massiv?) or another functional language (Futhark?). And a smart compiler really does have a shot at finding optimizations/fusions.

While I was at it, unrelated to the paper above, I think I made a working generic auto differentiable fold lens combinator. Pretty cool. mapAccumL is a hot tip.

For practical Haskell purposes, all of this is a little silly with the good Haskell AD packages around, the most prominent being


It is somewhat interesting to note the similarity of type forall s. Lens s appearing in the Matsuda and Wang approach to those those of the forall s. BVar s monad appearing in the backprop package. In this case I believe that the s type variable plays the same role it does in the ST monad, protecting a mutating Wengert tape state held in the monad, but I haven’t dug much into it. I don’t know enough about backprop to know what to make of this similarity.


The github repo with my playing around and stream of consciousness commentary is here

A Touch of Topological Quantum Computation in Haskell Pt. II: Automating Drudgery

Last time we built the basic pieces we need to describe anyons in Haskell. Anyon models describe interesting physical systems where a set of particles (Tau and Id in our case) have certain splitting rules and peculiar quantum properties. The existence of anyons in a system are the core physics necessary to support topological quantum computation. In topological quantum computing, quantum gates are applied by braiding the anyons and measurements performed by fusing anyons together and seeing what particle comes out. Applying gates in this way has inherent error correcting properties.

The tree of particle production with particle labelled leaves picks a basis (think the collection \{\hat{x}, \hat{y}, \hat{z}\} ) for the anyon quantum vector space. An individual basis vector (think \hat{x} ) from this basis is specified by labelling the internal edges of the tree. We built a Haskell data type for a basic free vector space and functions for the basic R-moves for braiding two anyons and reassociating the tree into a new basis with F-moves. In addition, you can move around your focus within the tree by using the function lmap and rmap. The github repo with that and what follows below is here.

Pain Points

We’ve built the atomic operations we need, but they work very locally and are quite manual. You can apply many lmap and rmap to zoom in to the leaves you actually wish to braid, and you can manually perform all the F-moves necessary to bring nodes under the same parent, but it will be rather painful.

The standard paper-and-pencil graphical notation for anyons is really awesome. You get to draw little knotty squiggles to calculate. It does not feel as laborious. The human eye and hand are great at applying a sequence of reasonably optimal moves to untangle the diagram efficiently. Our eye can take the whole thing in and our hand can zip around anywhere.

To try and bridge this gap, we need to build functions that work in some reasonable way on the global anyon tree and that automate simple tasks.

A Couple Useful Functions

Our first useful operation is pullLeftLeaf. This operation will rearrange the tree using F-moves to get the leftmost leaf associated all the way to the root. The leftmost leaf will then have the root as a parent.

Because the tree structure is in the FibTree a b data type, we need the tuple tree type of the pulled tree. This is a slightly non-trivial type computation.

In order to do this, we’ll use a bit of typelevel programming. If this is strange and alarming stuff for you, don’t sweat it too much. I am not the most elegant user of these techniques, but I hope that alongside my prose description you can get the gist of what we’re going for.

(Sandy Maguire has a new book on typelevel programming in Haskell out. Good stuff. Support your fellow Haskeller and toss him some buckos.)

The resulting tree type b is an easily computable function of the starting tree type a. That is what the “functional dependency” notation | a -> b in the typeclass definition tells the compiler.

The first 4 instances are base cases. If you’re all the way at the leaf, you basically want to do nothing. pure is the function that injects the classical tree description into a quantum state vector with coefficient 1.

The meat is in the last instance. In the case that the tree type matches ((a,b),c), we recursively call PullLeftLeaf on (a,b) which returns a new result (a',b'). Because of the recursion, this a' is the leftmost leaf. We can then construct the return type by doing a single reassociation step. The notation ~ forces two types to unify. We can use this conceptually as an assignment statement at the type level. This is very useful for building intermediate names for large expressions, as assert statements to ensure the types are as expected, and also occasionally to force unification of previously unknown types. It’s an interesting operator for sure.

The recursion at the type level is completely reflected in the actual function definition. We focus on the piece (a,b) inside t by using lmap. We do a recursive call to pullLeftLeaf, and finally fmove' performs the final reassociation move. It is all rather verbose, but straightforward I hope.

You can also build a completely similar PullRightLeaf.

A Canonical Right Associated Basis

One common way of dealing with larger trees is to pick a canonical basis of fully right associated trees. The fully right associated tree is a list-like structure. Its uniformity makes it easier to work with.

By recursively applying pullLeftLeaf, we can fully right associate any tree.

This looks quite similar to the implementation of pullLeftLeaf. It doesn’t actually have much logic to it. We apply pullLeftLeaf, then we recursively apply rightAssoc in the right branch of the tree.

B-Moves: Braiding in the Right Associated Basis

Now we have the means to convert any structure to it’s right associated canonical basis. In this basis, one can apply braiding to neighboring anyons using B-moves, which can be derived from the braiding R-moves and F-moves.

The B-move applies one F-move so that the two neighboring leaves share a parent, uses the regular braiding R-move, then applies the inverse F-move to return back to the canonical basis. Similarly, bmove'  is the same thing except applies the under braiding braid' rather that the over braiding braid.

(Image Source : Preskill’s notes)

Indexing to Leaves

We also may desire just specifying the integer index of where we wish to perform a braid. This can be achieved with another typeclass for iterated rmaping. When the tree is in canonical form, this will enable us to braid two neighboring leaves by an integer index. This index has to be a typelevel number because the output type depends on it.

In fact there is quite a bit of type computation. Given a total tree type s and an index n this function will zoom into the subpart a of the tree at which we want to apply our function. The subpart a is replaced by b, and then the tree is reconstructed into t. t is s with the subpart a mapped into b.  I have intentionally made this reminiscent of the type variables of the lens type Lens s t a b .

This looks much noisier that it has to because we need to work around some of the unfortunate realities of using the typeclass system to compute types. We can’t just match on the number n in order to pick which instance to use because the patterns 0 and n are overlapping. The pattern n can match the number 0 if n ~ 0. The pattern matching in the type instance is not quite the same as the regular Haskell pattern matching we use to define functions. The order of the definitions does not matter, so you can’t have default cases. The patterns you use cannot be unifiable. In order to fix this, we make the condition if n is greater than 0 an explicit type variable gte. Now the different cases cannot unify. It is a very common trick to need a variable representing some branching condition.

For later convenience, we define rmapN which let’s us not need to supply the necessary comparison type gte.

Parentifying Leaves Lazily

While it is convenient to describe anyon computations in a canonical basis, it can be quite inefficient. Converting an arbitrary  anyon tree into the standard basis will often result in a dense vector. A natural thing to do for the sake of economy is only do reassociation on demand.

The algorithm for braiding two neighboring leaves is pretty straightforward. We need to reassociate these leaves so that they have the same parent. First we need the ability to map into the least common ancestor of the two leaves. To reassociate these two leaves to have a common parent we pullrightLeaf the left subtree and then pullLeftLeaf the left subtree. Finally, there is a bit extra bit of shuffling to actually get them to be neighbors.

As a first piece, we need a type level function to count the number of leaves in a tree. In this case, I am inclined to use type families rather than multi parameter typeclasses as before, since I don’t need value level stuff coming along for the ride.

Next, we make a typeclass for mapping into the least common ancestor position.

We find the least common ancestor position by doing a binary search on the size of the left subtrees at each node. Once the size of the left subtree equals n, we’ve found the common ancestor of leaf n and leaf n+1.

Again, this LCAMap typeclass has a typelevel argument gte that directs it which direction to go down the tree.

The Twiddle typeclass will perform some final cleanup after we’ve done all the leaf pulling. At that point, the leaves still do not have the same parent. They are somewhere between 0 and 2 F-moves off depending on whether the left or right subtrees may be just a leaf or larger trees. twiddle is not a recursive function.

Putting this all together we get the nmap function that can apply a function after parentifying two leaves. By far the hardest part is writing out that type signature.

Usage Example

Here’s some simple usage:

Note that rmapN is 0-indexed but nmap is 1-indexed. This is somewhat horrifying, but that is what was natural in the implementation.

Here is a more extended example showing how to fuse some particles.

I started with the tree at the top and traversed downward implementing each braid and fusion. Implicitly all the particles shown in the diagram are Tau particles. The indices refer to particle position, not to the particles “identity” as you would trace it by eye on the page. Since these are identical quantum  particles, the particles don’t have identity as we classically think of it anyhow.

The particle pairs are indexed by the number on the left particle. First braid 1 over 2, then 2 over 3, fuse 1 and 2, braid 2 under 3, fuse 2 and 3, and then fuse 1 and 2. I got an amplitude for the process of -0.618, corresponding to a probability of 0.382. I would give myself 70% confidence that I implemented all my signs and conventions correctly. The hexagon and pentagon equations from last time being correct gives me some peace of mind.

Syntax could use a little spit polish, but it is usable. With some readjustment, one could use the Haskell do notation removing the need for explicit >>=.

Next Time

Anyons are often described in categorical terminology. Haskell has a category culture as well. Let’s explore how those mix!

Compiling to Categories 3: A Bit Cuter

Ordinary Haskell functions form a cartesian closed category. Category means you can compose functions. Cartesian basically means you can construct and deconstruct tuples and Closed means that you have first class functions you can pass around.

Conal Elliott’s Compiling to Categories is a paradigm for reinterpreting ordinary functions as the equivalent in other categories. At an abstract level, I think you could describe it as a mechanism to build certain natural law abiding Functors from Hask to other categories. It’s another way to write things once and have them run many ways, like polymorphism or generic programming. The ordinary function syntax is human friendly compared to writing raw categorical definitions, which look roughly like point-free programming (no named variables). In addition, by embedding it as a compiler pass within GHC, he gets to leverage existing GHC optimizations as optimizations for other categories. Interesting alternative categories include the category of graphs, circuits, and automatically differentiable functions. You can find his implementation here

I’ve felt hesitance at using a GHC plugin plus I kind of want to do it in a way I understand, so I’ve done different versions of this using relatively normal Haskell (no template haskell, no core passes, but a butt ton of hackery).

The first used explicit tags. Give them to the function and see where they come out. That is one way to probe a simple tuple rearrangement function.

The second version worked almost entirely at the typelevel. It worked on the observation that a completely polymorphic tuple type signature completely specifies the implementation. You don’t have to look at the term level at all. It unified the polymorphic values in the input with a typelevel number indexed Tag type. Then it searched through the input type tree to find the piece it needed. I did end up passing stuff in to the term level because I could use this mechanism to embed categorical literals. The typeclass hackery I used to achieve this all was heinous.

I realized today another way related to both that is much simpler and fairly direct. It has some pleasing aesthetic properties and some bad ones. The typeclass hackery is much reduced and the whole thing fits on a screen, so I’m pleased.

Here are the basic categorical definitions. FreeCat is useful for inspection in GHCi of what comes out of toCCC.

And here is the basic toCCC implementation


Here is some example usage


What we do is generate a tuple to give to your function. The function is assumed to be polymorphic again but now not necessarily totally polymorphic (this is important because Num typeclass usage will unify variables). Once we hit a leaf of the input tuple, we place the categorical morphism that would extract that piece from the input. For example for the input type (a,(b,c)) we pass it the value (fstC ,(fstC . sndC, sndC . sndC )). Detecting when we are at a leaf again requires somehow detecting a polymorphic location, which is a weird thing to do. We use the Incoherent IsTup instance from last time to do this. It is close to being safe, because we immediately unify the polymorphic variable with a categorial type, so if something goes awry, a type error should result. We could make it more ironclad by unifying immediately to something that contains the extractor and a user inaccessible type.

We apply the function to this input. Now the output is a tuple tree of morphisms. We recursively traverse down this tree with a fanC for every tuple. This all barely requires any typelevel hackery. The typelevel stuff that is there is just so that I can traverse down tuple trees basically.

One benefit is that we can now use some ordinary typeclasses. We can make a simple implementation of Num for (k z a) like how we would make it for (z -> a). This let’s us use the regular (+) and (*)  operators for example.

What is not good is the performance. As it is, the implementation takes many global duplication of the input to create all the fanCs. In many categories, this is very wasteful.This may be a fixable problem, either via passing in more sophisticated objects that just the bare extraction morphisms to to input (CPS-ified? Path Lists?) or via the GHC rewrite rules mechanism. I have started to attempt that, but have not been successful in getting any of my rewrite rules to fire yet, because I have no idea what I’m doing. If anyone could give me some advice, I’d be much obliged. You can check that out here. For more on rewrite rules, check out the GHC user manual and this excellent tutorial by Mark Karpov here.

Another possibility is to convert to FreeCat, write regular Haskell function optimization passes over the FreeCat AST and then interpret it. This adds interpretation overhead, which may or may not be acceptable depending on your use case. It would probably not be appropriate for automatically differentiable functions, but may be for outputting circuits or graphs.


Another problem is dealing with boolean operations. The booleans operators and comparison operators are not sufficiently polymorphic in the Prelude. We could define new operators that work as drop in replacements in the original context, but I don’t really have the ability to overload the originals. It is tough because if we do things like this, it feels like we’re really kind of building a DSL more than we are compiling to categories. We need to write our functions with the DSL in mind and can’t just import and use some function that had no idea about the compiling to categories stuff.

I should probably just be using Conal’s concat project. This is all a little silly.