Fix spelling mistakes
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doc/appendix.tex
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37
doc/appendix.tex
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\lstset{basicstyle=\footnotesize\ttfamily,breaklines=true,breakpages=true}
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\def\fileps
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{ ../src/Cat.agda
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, ../src/Cat/Categories/Cat.agda
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, ../src/Cat/Categories/Cube.agda
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, ../src/Cat/Categories/CwF.agda
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, ../src/Cat/Categories/Fam.agda
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, ../src/Cat/Categories/Free.agda
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, ../src/Cat/Categories/Fun.agda
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, ../src/Cat/Categories/Rel.agda
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, ../src/Cat/Categories/Sets.agda
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, ../src/Cat/Category.agda
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, ../src/Cat/Category/CartesianClosed.agda
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, ../src/Cat/Category/Exponential.agda
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, ../src/Cat/Category/Functor.agda
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, ../src/Cat/Category/Monad.agda
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, ../src/Cat/Category/Monad/Kleisli.agda
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, ../src/Cat/Category/Monad/Monoidal.agda
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, ../src/Cat/Category/Monad/Voevodsky.agda
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, ../src/Cat/Category/Monoid.agda
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, ../src/Cat/Category/NaturalTransformation.agda
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, ../src/Cat/Category/Product.agda
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, ../src/Cat/Category/Yoneda.agda
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, ../src/Cat/Equivalence.agda
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, ../src/Cat/Prelude.agda
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}
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\foreach \filep in \fileps {
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\chapter{\filep}
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%% \begin{figure}[htpb]
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\lstinputlisting{\filep}
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%% \caption{Source code for \texttt{\filep}}
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%% \label{fig:\filep}
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%% \end{figure}
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}
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%% \lstset{framextopmargin=50pt}
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%% \lstinputlisting{../../src/Cat.agda}
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@ -40,6 +40,7 @@
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\newcommand*{\subtitle}[1]{\gdef\@subtitle{#1}}
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\newcommand*{\subtitle}[1]{\gdef\@subtitle{#1}}
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\newgeometry{top=3cm, bottom=3cm,left=2.25 cm, right=2.25cm}
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\newgeometry{top=3cm, bottom=3cm,left=2.25 cm, right=2.25cm}
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\begingroup
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\begingroup
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\thispagestyle{empty}
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\usepackage{noto}
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\usepackage{noto}
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\fontseries{sb}
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\fontseries{sb}
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%% \fontfamily{noto}\selectfont
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%% \fontfamily{noto}\selectfont
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3
doc/conclusion.tex
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3
doc/conclusion.tex
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\chapter{Conclusion}
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\TODO{\ldots}
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@ -56,7 +56,7 @@ the real numbers from $0$ to $1$. $P$ is a family of types over the index set
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$I$. I will sometimes refer to $P$ as the ``path-space'' of some path $p \tp
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$I$. I will sometimes refer to $P$ as the ``path-space'' of some path $p \tp
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\Path\ P\ a\ b$. By this token $P\ 0$ then corresponds to the type at the
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\Path\ P\ a\ b$. By this token $P\ 0$ then corresponds to the type at the
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left-endpoint and $P\ 1$ as the type at the right-endpoint. The type is called
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left-endpoint and $P\ 1$ as the type at the right-endpoint. The type is called
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$\Path$ because it is connected with paths in homotopy thoery. The intuition
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$\Path$ because it is connected with paths in homotopy theory. The intuition
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behind this is that $\Path$ describes paths in $\MCU$ -- i.e. between types. For
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behind this is that $\Path$ describes paths in $\MCU$ -- i.e. between types. For
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a path $p$ for the point $p\ i$ the index $i$ describes how far along the path
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a path $p$ for the point $p\ i$ the index $i$ describes how far along the path
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one has moved. An inhabitant of $\Path\ P\ a_0\ a_1$ is a (dependent-)
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one has moved. An inhabitant of $\Path\ P\ a_0\ a_1$ is a (dependent-)
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@ -155,7 +155,7 @@ there is the notion of sets from set-theory, in Agda types are denoted
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\texttt{Set}. I will use it consistently to refer to a type $A$ as a set exactly
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\texttt{Set}. I will use it consistently to refer to a type $A$ as a set exactly
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if $\isSet\ A$ is inhabited.
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if $\isSet\ A$ is inhabited.
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The next step in the hierarchy is, as you might've guessed, the type:
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The next step in the hierarchy is, as the reader might've guessed, the type:
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%
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%
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\begin{equation}
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\begin{equation}
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\begin{alignat}{2}
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\begin{alignat}{2}
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@ -166,7 +166,7 @@ The next step in the hierarchy is, as you might've guessed, the type:
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%
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%
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And so it continues. In fact we can generalize this family of types by indexing
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And so it continues. In fact we can generalize this family of types by indexing
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them with a natural number. For historical reasons, though, the bottom of the
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them with a natural number. For historical reasons, though, the bottom of the
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hierarchy, the contractible tyes, is said to be a \nomen{-2-type}, propositions
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hierarchy, the contractible types, is said to be a \nomen{-2-type}, propositions
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are \nomen{-1-types}, (homotopical) sets are \nomen{0-types} and so on\ldots
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are \nomen{-1-types}, (homotopical) sets are \nomen{0-types} and so on\ldots
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Just as with paths, homotopical sets are not at the center of focus for this
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Just as with paths, homotopical sets are not at the center of focus for this
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@ -181,12 +181,12 @@ Proposition: For any homotopic level $n$ this is a mere proposition.
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%
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%
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\section{A few lemmas}
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\section{A few lemmas}
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Rather than getting into the nitty-gritty details of Agda I venture to take a
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Rather than getting into the nitty-gritty details of Agda I venture to take a
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more ``combinators-based'' approach. That is, I will use theorems about paths
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more ``combinator-based'' approach. That is, I will use theorems about paths
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already that have already been formalized. Specifically the results come from
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already that have already been formalized. Specifically the results come from
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the Agda library \texttt{cubical} (\TODO{Cite}). I have used a handful of
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the Agda library \texttt{cubical} (\TODO{Cite}). I have used a handful of
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results from this library as well as contributed a few lemmas myself.\footnote{The module \texttt{Cat.Prelude} lists the upstream dependencies. As well my contribution to \texttt{cubical} can be found in the git logs \TODO{Cite}.}
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results from this library as well as contributed a few lemmas myself.\footnote{The module \texttt{Cat.Prelude} lists the upstream dependencies. As well my contribution to \texttt{cubical} can be found in the git logs \TODO{Cite}.}
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These theorems are all purely related to homotopy theory and cubical agda and as
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These theorems are all purely related to homotopy theory and cubical Agda and as
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such not specific to the formalization of Category Theory. I will present a few
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such not specific to the formalization of Category Theory. I will present a few
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of these theorems here, as they will be used later in chapter
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of these theorems here, as they will be used later in chapter
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\ref{ch:implementation} throughout.
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\ref{ch:implementation} throughout.
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@ -195,7 +195,7 @@ of these theorems here, as they will be used later in chapter
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\label{sec:pathJ}
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\label{sec:pathJ}
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The induction principle for paths intuitively gives us a way to reason about a
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The induction principle for paths intuitively gives us a way to reason about a
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type-family indexed by a path by only considering if said path is $\refl$ (the
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type-family indexed by a path by only considering if said path is $\refl$ (the
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``base-case''). For \emph{based path induction}, that equaility is \emph{based}
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``base-case''). For \emph{based path induction}, that equality is \emph{based}
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at some element $a \tp A$.
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at some element $a \tp A$.
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Let a type $A \tp \MCU$ and an element of the type $a \tp A$ be given. $a$ is said to be the base of the induction. Given a family of types:
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Let a type $A \tp \MCU$ and an element of the type $a \tp A$ be given. $a$ is said to be the base of the induction. Given a family of types:
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74
doc/discussion.tex
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74
doc/discussion.tex
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\chapter{Perspectives}
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\section{Discussion}
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In the previous chapter the practical aspects of proving things in Cubical Agda
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were highlighted. I also demonstrated the usefulness of separating ``laws'' from
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``data''. One of the reasons for this is that dependencies within types can lead
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to very complicated goals. One technique for alleviating this was to prove that
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certain types are mere propositions.
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\subsection{Computational properties}
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Another aspect (\TODO{That I actually didn't highlight very well in the previous
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chapter}) is the computational nature of paths. Say we have formalized this
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common result about monads:
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\TODO{Some equation\ldots}
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By transporting this to the Kleisli formulation we get a result that we can use
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to compute with. This is particularly useful because the Kleisli formulation
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will be more familiar to programmers e.g. those coming from a background in
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Haskell. Whereas the theory usually talks about monoidal monads.
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\TODO{Mention that with postulates we cannot do this}
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\subsection{Reusability of proofs}
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The previous example also illustrate how univalence unifies two otherwise
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disparate areas: The category-theoretic study of monads; and monads as in
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functional programming. Univalence thus allows one to reuse proofs. You could
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say that univalence gives the developer two proofs for the price of one.
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The introduction (section \ref{sec:context}) mentioned an often
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employed-technique for enabling extensional equalities is to use the
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setoid-interpretation. Nowhere in this formalization has this been necessary,
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$\Path$ has been used globally in the project as propositional equality. One
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interesting place where this becomes apparent is in interfacing with the Agda
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standard library. Multiple definitions in the Agda standard library have been
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designed with the setoid-interpretation in mind. E.g. the notion of ``unique
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existential'' is indexed by a relation that should play the role of
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propositional equality. Likewise for equivalence relations, they are indexed,
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not only by the actual equivalence relation, but also by another relation that
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serve as propositional equality.
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%% Unfortunately we cannot use the definition of equivalences found in the
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%% standard library to do equational reasoning directly. The reason for this is
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%% that the equivalence relation defined there must be a homogenous relation,
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%% but paths are heterogeneous relations.
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In the formalization at present a significant amount of energy has been put
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towards proving things that would not have been needed in classical Agda. The
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proofs that some given type is a proposition were provided as a strategy to
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simplify some otherwise very complicated proofs (e.g.
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\ref{eq:proof-prop-IsPreCategory} and \label{eq:productPath}). Often these
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proofs would not be this complicated. If the J-rule holds definitionally the
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proof-assistant can help simplify these goals considerably. The lack of the
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J-rule has a significant impact on the complexity of these kinds of proofs.
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\TODO{Universe levels.}
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\section{Future work}
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\subsection{Agda \texttt{Prop}}
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Jesper Cockx' work extending the universe-level-laws for Agda and the
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\texttt{Prop}-type.
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\subsection{Compiling Cubical Agda}
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\label{sec:compiling-cubical-agda}
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Compilation of program written in Cubical Agda is currently not supported. One
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issue here is that the backends does not provide an implementation for the
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cubical primitives (such as the path-type). This means that even though the
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path-type gives us a computational interpretation of functional extensionality,
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univalence, transport, etc., we do not have a way of actually using this to
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compile our programs that use these primitives. It would be interesting to see
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practical applications of this. The path between monads that this library
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exposes could provide one particularly interesting case-study.
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\subsection{Higher inductive types}
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This library has not explored the usefulness of higher inductive types in the
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context of Category Theory.
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\label{ch:implementation}
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\label{ch:implementation}
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This implementation formalizes the following concepts:
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This implementation formalizes the following concepts:
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%
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%
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\begin{itemize}
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\begin{enumerate}[i.]
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\item Core categorical concepts
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\item Categories
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\subitem Categories
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\item Functors
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\subitem Functors
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\item Products
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\subitem Products
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\item Exponentials
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\subitem Exponentials
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\item Cartesian closed categories
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\subitem Cartesian closed categories
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\item Natural transformations
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\subitem Natural transformations
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\item Yoneda embedding
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\subitem Yoneda embedding
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\item Monads
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\subitem Monads
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\item Categories
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\subsubitem Monoidal monads
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\begin{enumerate}[i.]
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\subsubitem Kleisli monads
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\item Opposite category
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\subsubitem Voevodsky's construction
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\item Category of sets
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\item Category of \ldots
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\item ``Pair category''
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\subitem Homotopy sets
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\end{enumerate}
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\subitem Categories -- only data-part
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\end{enumerate}
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\subitem Relations -- only data-part
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\subitem Functors -- only as a precategory
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\subitem Free category
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\end{itemize}
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%
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%
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Since it is useful to distinguish between types with more or less (homotopical)
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Furthermore the following items have been partly formalized:
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structure I have followed the following design-principle: I have split concepts
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%
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up into things that represent ``data'' and ``laws'' about this data. The idea is
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\begin{enumerate}[i.]
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that we can provide a proof that the laws are mere propositions. As an example a
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\item The (higher) category of categories.
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category is defined to have two members: `raw` which is a collection of the data
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\item Category of relations
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and `isCategory` which asserts some laws about that data.
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\item Category of functors and natural transformations -- only as a precategory
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\item Free category
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\item Monoidal objects
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\item Monoidal categories
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\end{enumerate}
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%
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As well as a range of various results about these. E.g. I have shown that the
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category of sets has products. In the following I aim to demonstrate some of the
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techniques employed in this formalization and in the interest of brevity I will
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not detail all the things I have formalized. In stead, I have selected a parts
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of this formalization that highlight some interesting proof techniques relevant
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to doing proofs in Cubical Agda.
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One such technique that is pervasive to this formalization is the idea of
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distinguishing types with more or less homotopical structure. To do this I have
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followed the following design-principle: I have split concepts up into things
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that represent ``data'' and ``laws'' about this data. The idea is that we can
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provide a proof that the laws are mere propositions. As an example a category is
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defined to have two members: `raw` which is a collection of the data and
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`isCategory` which asserts some laws about that data.
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This allows me to reason about things in a more ``standard mathematical way'',
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This allows me to reason about things in a more ``standard mathematical way'',
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where one can reason about two categories by simply focusing on the data. This
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where one can reason about two categories by simply focusing on the data. This
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is acheived by creating a function embodying the ``equality principle'' for a
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is achieved by creating a function embodying the ``equality principle'' for a
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given type.
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given type.
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\section{Categories}
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\section{Categories}
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@ -45,7 +60,7 @@ on the topic. We, however, impose one further requirement on what it means to be
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a category, namely that the type of arrows form a set.
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a category, namely that the type of arrows form a set.
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Such categories are called \nomen{1-categories}. It's possible to relax this
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Such categories are called \nomen{1-categories}. It's possible to relax this
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requirement. This would lead to the notion of higher categorier (\cite[p.
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requirement. This would lead to the notion of higher categories (\cite[p.
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307]{hott-2013}). For the purpose of this project, however, this report will
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307]{hott-2013}). For the purpose of this project, however, this report will
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restrict itself to 1-categories. Making based on higher categories would be a
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restrict itself to 1-categories. Making based on higher categories would be a
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very natural possible extension of this work.
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very natural possible extension of this work.
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@ -63,7 +78,7 @@ definitions: If we let $p$ be a witness to the identity law, which formally is:
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\end{equation}
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\end{equation}
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%
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%
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Then we can construct the identity isomorphism $\var{idIso} \tp \identity,
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Then we can construct the identity isomorphism $\var{idIso} \tp \identity,
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\identity, p \tp A \approxeq A$ for any obejct $A$. Here $\approxeq$ denotes
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\identity, p \tp A \approxeq A$ for any object $A$. Here $\approxeq$ denotes
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isomorphism on objects (whereas $\cong$ denotes isomorphism of types). This will
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isomorphism on objects (whereas $\cong$ denotes isomorphism of types). This will
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be elaborated further on in sections \ref{sec:equiv} and \ref{sec:univalence}.
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be elaborated further on in sections \ref{sec:equiv} and \ref{sec:univalence}.
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Moreover, due to substitution for paths we can construct an isomorphism from
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Moreover, due to substitution for paths we can construct an isomorphism from
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@ -121,7 +136,7 @@ f \lll \identity ≡ f
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\end{align}
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\end{align}
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%
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%
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$\lll$ denotes arrow composition (right-to-left), and reverse function
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$\lll$ denotes arrow composition (right-to-left), and reverse function
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composition (left-to-right, diagramatic order) is denoted $\rrr$. The objects
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composition (left-to-right, diagrammatic order) is denoted $\rrr$. The objects
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($A$, $B$ and $C$) and arrow ($f$, $g$, $h$) are implicitly universally
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($A$, $B$ and $C$) and arrow ($f$, $g$, $h$) are implicitly universally
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quantified.
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quantified.
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@ -160,25 +175,32 @@ us the two obligations: $\isProp\ (\id \comp f \equiv f)$ and $\isProp\ (f \comp
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set.
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set.
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This example illustrates nicely how we can use these combinators to reason about
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This example illustrates nicely how we can use these combinators to reason about
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`canonical' types like $\sum$ and $\prod$. Similiar combinators can be defined
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`canonical' types like $\sum$ and $\prod$. Similar combinators can be defined
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at the other homotopic levels. These combinators are however not applicable in
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at the other homotopic levels. These combinators are however not applicable in
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situations where we want to reason about other types - e.g. types we've defined
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situations where we want to reason about other types - e.g. types we've defined
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ourselves. For instance, after we've proven that all the projections of
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ourselves. For instance, after we've proven that all the projections of
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pre-categories are propositions, then we would like to bundle this up to show
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pre-categories are propositions, then we would like to bundle this up to show
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that the type of pre-categories is also a proposition. Since pre-categories are
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that the type of pre-categories is also a proposition. Formally:
|
||||||
not formulated with a chain of sigma-types we wont have any combinators
|
|
||||||
available to help us here. In stead we'll have to use the path-type directly.
|
|
||||||
|
|
||||||
What we want to prove is:
|
|
||||||
%
|
%
|
||||||
\begin{equation}
|
\begin{equation}
|
||||||
\label{eq:propIsPreCategory}
|
\label{eq:propIsPreCategory}
|
||||||
\isProp\ \IsPreCategory
|
\isProp\ \IsPreCategory
|
||||||
\end{equation}
|
\end{equation}
|
||||||
%
|
%
|
||||||
%% \begin{proof}
|
Where The definition of $\IsPreCategory$ is the triple:
|
||||||
%
|
%
|
||||||
This is judgmentally the same as
|
\begin{align*}
|
||||||
|
\var{isAssociative} & \tp \var{IsAssociative}\\
|
||||||
|
\var{isIdentity} & \tp \var{IsIdentity}\\
|
||||||
|
\var{arrowsAreSets} & \tp \var{ArrowsAreSets}
|
||||||
|
\end{align*}
|
||||||
|
%
|
||||||
|
Each corresponding to the first three laws for categories. Note that since
|
||||||
|
$\IsPreCategory$ is not formulated with a chain of sigma-types we wont have any
|
||||||
|
combinators available to help us here. In stead we'll have to use the path-type
|
||||||
|
directly.
|
||||||
|
|
||||||
|
\ref{eq:propIsPreCategory} is judgmentally the same as
|
||||||
%
|
%
|
||||||
$$
|
$$
|
||||||
\prod_{a\ b \tp \IsPreCategory} a \equiv b
|
\prod_{a\ b \tp \IsPreCategory} a \equiv b
|
||||||
|
@ -202,12 +224,13 @@ So to give the continuous function $I \to \IsPreCategory$, which is our goal, we
|
||||||
introduce $i \tp I$ and proceed by constructing an element of $\IsPreCategory$
|
introduce $i \tp I$ and proceed by constructing an element of $\IsPreCategory$
|
||||||
by using the fact that all the projections are propositions to generate paths
|
by using the fact that all the projections are propositions to generate paths
|
||||||
between all projections. Once we have such a path e.g. $p \tp a.\isIdentity
|
between all projections. Once we have such a path e.g. $p \tp a.\isIdentity
|
||||||
\equiv b.\isIdentity$ we can elimiate it with $i$ and thus obtain $p\ i \tp
|
\equiv b.\isIdentity$ we can eliminate it with $i$ and thus obtain $p\ i \tp
|
||||||
(p\ i).\isIdentity$. This element satisfies exactly that it corresponds to the
|
(p\ i).\isIdentity$. This element satisfies exactly that it corresponds to the
|
||||||
corresponding projections at either endpoint. Thus the element we construct at
|
corresponding projections at either endpoint. Thus the element we construct at
|
||||||
$i$ becomes the triple:
|
$i$ becomes the triple:
|
||||||
%
|
%
|
||||||
\begin{equation}
|
\begin{equation}
|
||||||
|
\label{eq:proof-prop-IsPreCategory}
|
||||||
\begin{aligned}
|
\begin{aligned}
|
||||||
& \var{propIsAssociative} && a.\var{isAssociative}\
|
& \var{propIsAssociative} && a.\var{isAssociative}\
|
||||||
&& b.\var{isAssociative} && i \\
|
&& b.\var{isAssociative} && i \\
|
||||||
|
@ -224,7 +247,7 @@ It is worth noting that proving this theorem with the regular inductive equality
|
||||||
type would already not be possible, since we at least need extensionality (the
|
type would already not be possible, since we at least need extensionality (the
|
||||||
projections are all $\prod$-types). Assuming we had functional extensionality
|
projections are all $\prod$-types). Assuming we had functional extensionality
|
||||||
available to us as an axiom, we would use functional extensionality (in
|
available to us as an axiom, we would use functional extensionality (in
|
||||||
reverse?) to retreive the equalities in $a$ and $b$, pattern-match on them to
|
reverse?) to retrieve the equalities in $a$ and $b$, pattern-match on them to
|
||||||
see that they are both $\var{refl}$ and then close the proof with $\var{refl}$.
|
see that they are both $\var{refl}$ and then close the proof with $\var{refl}$.
|
||||||
Of course this theorem is not so interesting in the setting of ITT since we know
|
Of course this theorem is not so interesting in the setting of ITT since we know
|
||||||
a priori that equality proofs are unique.
|
a priori that equality proofs are unique.
|
||||||
|
@ -234,7 +257,7 @@ instance, when we want to show that $\IsCategory$ is a mere proposition.
|
||||||
$\IsCategory$ is a record with two fields, a witness to being a pre-category and
|
$\IsCategory$ is a record with two fields, a witness to being a pre-category and
|
||||||
the univalence condition. Recall that the univalence condition is indexed by the
|
the univalence condition. Recall that the univalence condition is indexed by the
|
||||||
identity-proof. So to follow the same recipe as above, let $a\ b \tp
|
identity-proof. So to follow the same recipe as above, let $a\ b \tp
|
||||||
\IsCategory$ be given, to show them equal, we now need to give two paths. One homogenous:
|
\IsCategory$ be given, to show them equal, we now need to give two paths. One homogeneous:
|
||||||
%
|
%
|
||||||
$$
|
$$
|
||||||
p \tp a.\isPreCategory \equiv b.\isPreCategory
|
p \tp a.\isPreCategory \equiv b.\isPreCategory
|
||||||
|
@ -249,9 +272,9 @@ $$
|
||||||
Which depends on the choice of $p$. The first of these we can provide since, as
|
Which depends on the choice of $p$. The first of these we can provide since, as
|
||||||
we have shown, $\IsPreCategory$ is a proposition. However, even though
|
we have shown, $\IsPreCategory$ is a proposition. However, even though
|
||||||
$\Univalent$ is also a proposition, we cannot use this directly to show the
|
$\Univalent$ is also a proposition, we cannot use this directly to show the
|
||||||
latter. This is becasue $\isProp$ talks about non-dependent paths. So we need to
|
latter. This is because $\isProp$ talks about non-dependent paths. So we need to
|
||||||
'promote' the result that univalence is a proposition to a heterogeneous path.
|
'promote' the result that univalence is a proposition to a heterogeneous path.
|
||||||
To this end we can use $\lemPropF$, which wasintroduced in \ref{sec:lemPropF}.
|
To this end we can use $\lemPropF$, which was introduced in \ref{sec:lemPropF}.
|
||||||
|
|
||||||
In this case $A = \var{IsIdentity}\ \identity$ and $B = \var{Univalent}$. We've
|
In this case $A = \var{IsIdentity}\ \identity$ and $B = \var{Univalent}$. We've
|
||||||
shown that being a category is a proposition, a result that holds for any choice
|
shown that being a category is a proposition, a result that holds for any choice
|
||||||
|
@ -267,7 +290,7 @@ And this finishes the proof that being-a-category is a mere proposition
|
||||||
(\ref{eq:propIsPreCategory}).
|
(\ref{eq:propIsPreCategory}).
|
||||||
|
|
||||||
When we have a proper category we can make precise the notion of ``identifying
|
When we have a proper category we can make precise the notion of ``identifying
|
||||||
isomorphic types'' \TODO{cite awodey here}. That is, we can construct the
|
isomorphic types'' \TODO{cite Awodey here}. That is, we can construct the
|
||||||
function:
|
function:
|
||||||
%
|
%
|
||||||
$$
|
$$
|
||||||
|
@ -322,7 +345,7 @@ The maps $\var{fromIso}$ and $\var{toIso}$ naturally extend to these maps:
|
||||||
\end{align}
|
\end{align}
|
||||||
%
|
%
|
||||||
Having this interface gives us both: a way to think rather abstractly about how
|
Having this interface gives us both: a way to think rather abstractly about how
|
||||||
to work with equivalences and a way to use ad-hoc definitions of equivalences.
|
to work with equivalences and a way to use ad hoc definitions of equivalences.
|
||||||
The specific instantiation of $\isEquiv$ as defined in \cite{cubical-agda} is:
|
The specific instantiation of $\isEquiv$ as defined in \cite{cubical-agda} is:
|
||||||
%
|
%
|
||||||
$$
|
$$
|
||||||
|
@ -394,7 +417,7 @@ univalence of all pre-category since morphisms in a category are not regular
|
||||||
functions -- in stead they can be thought of as a generalization hereof. The univalence criterion therefore is simply a way of restricting arrows
|
functions -- in stead they can be thought of as a generalization hereof. The univalence criterion therefore is simply a way of restricting arrows
|
||||||
to behave similarly to maps.
|
to behave similarly to maps.
|
||||||
|
|
||||||
I will now mention a few helpful thoerems that follow from univalence that will
|
I will now mention a few helpful theorems that follow from univalence that will
|
||||||
become useful later.
|
become useful later.
|
||||||
|
|
||||||
Obviously univalence gives us an isomorphism between $A \equiv B$ and $A
|
Obviously univalence gives us an isomorphism between $A \equiv B$ and $A
|
||||||
|
@ -427,7 +450,7 @@ then have the following two theorems:
|
||||||
\var{coe}\ p_{\var{cod}}\ f \equiv \iota \lll f
|
\var{coe}\ p_{\var{cod}}\ f \equiv \iota \lll f
|
||||||
\end{align}
|
\end{align}
|
||||||
%
|
%
|
||||||
I will give the proof of the first theorem here, the second one is analagous.
|
I will give the proof of the first theorem here, the second one is analogous.
|
||||||
%
|
%
|
||||||
\begin{align*}
|
\begin{align*}
|
||||||
\var{coe}\ p_{\var{dom}}\ f
|
\var{coe}\ p_{\var{dom}}\ f
|
||||||
|
@ -467,11 +490,11 @@ The base-case therefore becomes:
|
||||||
& \equiv f \lll \inv{(\var{idToIso}\ \widetilde{\refl})}
|
& \equiv f \lll \inv{(\var{idToIso}\ \widetilde{\refl})}
|
||||||
\end{align*}
|
\end{align*}
|
||||||
%
|
%
|
||||||
The first step follows because reflixivity is a neutral element for coercions.
|
The first step follows because reflexivity is a neutral element for coercion.
|
||||||
The second step is the identity law in the category. The last step has to do
|
The second step is the identity law in the category. The last step has to do
|
||||||
with the fact that $\var{idToIso}$ is constructed by substituting according to
|
with the fact that $\var{idToIso}$ is constructed by substituting according to
|
||||||
the supplied path and since reflexivity is also the neutral element for
|
the supplied path and since reflexivity is also the neutral element for
|
||||||
substuitutions we arrive at the desired expression. To close the
|
substitutions we arrive at the desired expression. To close the
|
||||||
based-path-induction we must supply the value ``at the other''. In this case
|
based-path-induction we must supply the value ``at the other''. In this case
|
||||||
this is simply $B \tp \Object$ and $p \tp A \equiv B$ which we have.
|
this is simply $B \tp \Object$ and $p \tp A \equiv B$ which we have.
|
||||||
|
|
||||||
|
@ -552,7 +575,7 @@ follows:
|
||||||
&& \text{$\var{shuffle}$ is an isomorphism} \\
|
&& \text{$\var{shuffle}$ is an isomorphism} \\
|
||||||
& \equiv
|
& \equiv
|
||||||
\identity
|
\identity
|
||||||
&& \text{$\isoToId$ is an ismorphism}
|
&& \text{$\isoToId$ is an isomorphism}
|
||||||
\end{align*}
|
\end{align*}
|
||||||
%
|
%
|
||||||
The other direction is analogous.
|
The other direction is analogous.
|
||||||
|
@ -639,7 +662,7 @@ lemma:
|
||||||
\end{align}
|
\end{align}
|
||||||
%
|
%
|
||||||
Which says that if two type-families are equivalent at all points, then pairs
|
Which says that if two type-families are equivalent at all points, then pairs
|
||||||
with identitical first components and these families as second components will
|
with identical first components and these families as second components will
|
||||||
also be equivalent. For our purposes $P \defeq \isEquiv\ A\ B$ and $Q \defeq
|
also be equivalent. For our purposes $P \defeq \isEquiv\ A\ B$ and $Q \defeq
|
||||||
\var{Isomorphism}$. So we must finally prove:
|
\var{Isomorphism}$. So we must finally prove:
|
||||||
%
|
%
|
||||||
|
@ -648,7 +671,7 @@ also be equivalent. For our purposes $P \defeq \isEquiv\ A\ B$ and $Q \defeq
|
||||||
\prod_{f \tp A \to B} \left( \isEquiv\ A\ B\ f \simeq \var{Isomorphism}\ f \right)
|
\prod_{f \tp A \to B} \left( \isEquiv\ A\ B\ f \simeq \var{Isomorphism}\ f \right)
|
||||||
\end{align}
|
\end{align}
|
||||||
|
|
||||||
First, lets proove \ref{eq:equivPropSig}: Let $propP \tp \prod_{a \tp A} \isProp (P\ a)$ and $x\;y \tp \sum_{a \tp A} P\ a$ be given. Because
|
First, lets prove \ref{eq:equivPropSig}: Let $propP \tp \prod_{a \tp A} \isProp (P\ a)$ and $x\;y \tp \sum_{a \tp A} P\ a$ be given. Because
|
||||||
of $\var{fromIsomorphism}$ it suffices to give an isomorphism between
|
of $\var{fromIsomorphism}$ it suffices to give an isomorphism between
|
||||||
$x \equiv y$ and $\fst\ x \equiv \fst\ y$:
|
$x \equiv y$ and $\fst\ x \equiv \fst\ y$:
|
||||||
%
|
%
|
||||||
|
@ -674,7 +697,7 @@ the two pairs:
|
||||||
\end{align*}
|
\end{align*}
|
||||||
%
|
%
|
||||||
The proof that these are indeed inverses has been omitted. \TODO{Do I really
|
The proof that these are indeed inverses has been omitted. \TODO{Do I really
|
||||||
want to ommit it?}\QED
|
want to omit it?}\QED
|
||||||
|
|
||||||
Now to prove \ref{eq:equivSig}: Let $e \tp \prod_{a \tp A} \left( P\ a \simeq
|
Now to prove \ref{eq:equivSig}: Let $e \tp \prod_{a \tp A} \left( P\ a \simeq
|
||||||
Q\ a \right)$ be given. To prove the equivalence, it suffices to give an
|
Q\ a \right)$ be given. To prove the equivalence, it suffices to give an
|
||||||
|
@ -839,14 +862,13 @@ arrows form a set. For instance, to prove associativity we must prove that
|
||||||
(\overline{h} \lll \overline{g}) \lll \overline{f}
|
(\overline{h} \lll \overline{g}) \lll \overline{f}
|
||||||
\end{align}
|
\end{align}
|
||||||
%
|
%
|
||||||
Herer $\lll$ refers to the `embellished' composition and $\overline{f}$,
|
Here $\lll$ refers to the `embellished' composition and $\overline{f}$,
|
||||||
$\overline{g}$ and $\overline{h}$ are triples consisting of arrows from the
|
$\overline{g}$ and $\overline{h}$ are triples consisting of arrows from the
|
||||||
underlying category ($f$, $g$ and $h$) and a pair of witnesses to
|
underlying category ($f$, $g$ and $h$) and a pair of witnesses to
|
||||||
\ref{eq:pairArrowLaw}.
|
\ref{eq:pairArrowLaw}.
|
||||||
%% Luckily those winesses are paths in the hom-set of the
|
%% Luckily those winesses are paths in the hom-set of the
|
||||||
%% underlying category which is a set, so these are mere propositions.
|
%% underlying category which is a set, so these are mere propositions.
|
||||||
The proof
|
The proof obligations is consists of two things. The first one is:
|
||||||
obligations is:
|
|
||||||
%
|
%
|
||||||
\begin{align}
|
\begin{align}
|
||||||
\label{eq:productAssocUnderlying}
|
\label{eq:productAssocUnderlying}
|
||||||
|
@ -855,15 +877,18 @@ h \lll (g \lll f)
|
||||||
(h \lll g) \lll f
|
(h \lll g) \lll f
|
||||||
\end{align}
|
\end{align}
|
||||||
%
|
%
|
||||||
Which is provable by \TODO{What?} and that the witness to \ref{eq:pairArrowLaw}
|
And the other proof obligation is that the witness to \ref{eq:pairArrowLaw} for
|
||||||
for the left-hand-side and the right-hand-side are the same. The type of this
|
the left-hand-side and the right-hand-side are the same.
|
||||||
goal is quite involved, and I will not write it out in full, but at present it
|
|
||||||
suffices to show the type of the path-space. Note that the arrows in
|
The proof of the first goal comes directly from the underlying category. The
|
||||||
\ref{eq:productAssoc} are arrows from $\mathcal{A} = (A , a_{\pairA} , a_{\pairB})$
|
type of the second goal is very complicated. I will not write it out in full
|
||||||
to $\mathcal{D} = (D , d_{\pairA} , d_{\pairB})$ where $a_{\pairA}$, $a_{\pairB}$,
|
here, but it suffices to show the type of the path-space. Note that the arrows
|
||||||
$d_{\pairA}$ and $d_{\pairB}$ are arrows in the underlying category. Given that $p$
|
in \ref{eq:productAssoc} are arrows from $\mathcal{A} = (A , a_{\pairA} ,
|
||||||
is the chosen proof of \ref{eq:productAssocUnderlying} we then have that the
|
a_{\pairB})$ to $\mathcal{D} = (D , d_{\pairA} , d_{\pairB})$ where
|
||||||
witness to \ref{eq:pairArrowLaw} vary over the type:
|
$a_{\pairA}$, $a_{\pairB}$, $d_{\pairA}$ and $d_{\pairB}$ are arrows in the
|
||||||
|
underlying category. Given that $p$ is the chosen proof of
|
||||||
|
\ref{eq:productAssocUnderlying} we then have that the witness to
|
||||||
|
\ref{eq:pairArrowLaw} vary over the type:
|
||||||
%
|
%
|
||||||
\begin{align}
|
\begin{align}
|
||||||
\label{eq:productPath}
|
\label{eq:productPath}
|
||||||
|
@ -979,7 +1004,7 @@ the implementation for the details).
|
||||||
\TODO{Super complicated}
|
\TODO{Super complicated}
|
||||||
|
|
||||||
\emph{Proposition} \ref{eq:univ-2} is isomorphic to \ref{eq:univ-3}: For this I
|
\emph{Proposition} \ref{eq:univ-2} is isomorphic to \ref{eq:univ-3}: For this I
|
||||||
will swho two corrolaries of \ref{eq:coeCod}: For an isomorphism $(\iota,
|
will show two corollaries of \ref{eq:coeCod}: For an isomorphism $(\iota,
|
||||||
\inv{\iota}, \var{inv}) \tp A \cong B$, arrows $f \tp \Arrow\ A\ X$, $g \tp
|
\inv{\iota}, \var{inv}) \tp A \cong B$, arrows $f \tp \Arrow\ A\ X$, $g \tp
|
||||||
\Arrow\ B\ X$ and a heterogeneous path between them, $q \tp \Path\ (\lambda i
|
\Arrow\ B\ X$ and a heterogeneous path between them, $q \tp \Path\ (\lambda i
|
||||||
\to p_{\var{dom}}\ i)\ f\ g$, where $p_{\var{dom}} \tp \Arrow\ A\ X \equiv
|
\to p_{\var{dom}}\ i)\ f\ g$, where $p_{\var{dom}} \tp \Arrow\ A\ X \equiv
|
||||||
|
@ -995,7 +1020,7 @@ g & \equiv f \lll \inv{\iota}
|
||||||
%
|
%
|
||||||
Proof: \TODO{\ldots}
|
Proof: \TODO{\ldots}
|
||||||
|
|
||||||
Now we can prove the equiavalence in the following way: Given $(f, \inv{f},
|
Now we can prove the equivalence in the following way: Given $(f, \inv{f},
|
||||||
\var{inv}_f) \tp X \cong Y$ and two heterogeneous paths
|
\var{inv}_f) \tp X \cong Y$ and two heterogeneous paths
|
||||||
%
|
%
|
||||||
\begin{align*}
|
\begin{align*}
|
||||||
|
@ -1025,7 +1050,7 @@ Which, since $f$ is an isomorphism and $p_{\mathcal{A}}$ (resp. $p_{\mathcal{B}}
|
||||||
is a path varying according to a path constructed from this isomorphism, this is
|
is a path varying according to a path constructed from this isomorphism, this is
|
||||||
exactly what \ref{eq:domain-twist-0} gives us.
|
exactly what \ref{eq:domain-twist-0} gives us.
|
||||||
%
|
%
|
||||||
The other direction is quite analagous. We choose $\inv{f}$ as the morphism and
|
The other direction is quite analogous. We choose $\inv{f}$ as the morphism and
|
||||||
prove that it satisfies \ref{eq:pairArrowLaw} with \ref{eq:domain-twist-1}.
|
prove that it satisfies \ref{eq:pairArrowLaw} with \ref{eq:domain-twist-1}.
|
||||||
|
|
||||||
We must now show that this choice of arrows indeed form an isomorphism. Our
|
We must now show that this choice of arrows indeed form an isomorphism. Our
|
||||||
|
@ -1099,7 +1124,7 @@ The proof of the first one is:
|
||||||
We have now constructed the maps between \ref{eq:univ-0} and \ref{eq:univ-1}. It
|
We have now constructed the maps between \ref{eq:univ-0} and \ref{eq:univ-1}. It
|
||||||
remains to show that they are inverses of each other. To cut a long story short,
|
remains to show that they are inverses of each other. To cut a long story short,
|
||||||
the proof uses the fact that isomorphism-of is propositional and that arrows (in
|
the proof uses the fact that isomorphism-of is propositional and that arrows (in
|
||||||
both categories) are sets. The reader is refered to the implementation for the
|
both categories) are sets. The reader is referred to the implementation for the
|
||||||
gory details.
|
gory details.
|
||||||
%
|
%
|
||||||
\subsection{Propositionality of products}
|
\subsection{Propositionality of products}
|
||||||
|
@ -1137,7 +1162,7 @@ conditions will be equal since $f$ is unique.
|
||||||
|
|
||||||
For the other direction we are now given a product $X, x_𝒜, x_ℬ$. Again this
|
For the other direction we are now given a product $X, x_𝒜, x_ℬ$. Again this
|
||||||
will be the terminal object. So now it remains that for any other object there
|
will be the terminal object. So now it remains that for any other object there
|
||||||
is aunique arrow from that object into $X, x_𝒜, x_ℬ$. Let $Y, y_𝒜, y_ℬ$ be
|
is a unique arrow from that object into $X, x_𝒜, x_ℬ$. Let $Y, y_𝒜, y_ℬ$ be
|
||||||
another object. As the arrow $\Arrow\ Y\ X$ we choose the product-arrow $y_𝒜 \x
|
another object. As the arrow $\Arrow\ Y\ X$ we choose the product-arrow $y_𝒜 \x
|
||||||
y_ℬ$. Since this is a product-arrow it satisfies \ref{eq:pairCondRev}. Let us
|
y_ℬ$. Since this is a product-arrow it satisfies \ref{eq:pairCondRev}. Let us
|
||||||
name the witness to this $\phi_{y_𝒜 \x y_ℬ}$. So we have picked as our center of
|
name the witness to this $\phi_{y_𝒜 \x y_ℬ}$. So we have picked as our center of
|
||||||
|
@ -1184,12 +1209,14 @@ That in any category:
|
||||||
\end{align}
|
\end{align}
|
||||||
%
|
%
|
||||||
\section{Monads}
|
\section{Monads}
|
||||||
In this section I show two formulations of monads and then show that they are
|
In this section I present two formulations of monads. The two representations
|
||||||
the same. The two representations are referred to as the monoidal- and kleisli-
|
are referred to as the monoidal- and Kleisli- representation respectively or
|
||||||
representation respectively.
|
simply monoidal monads and Kleisli monads for short. We then show that the two
|
||||||
|
formulations are equivalent, which due to univalence gives us a path between the
|
||||||
|
two types.
|
||||||
|
|
||||||
We shall let a category $\bC$ be given. In the remainder all objects and arrows
|
Let a category $\bC$ be given. In the remainder of this sections all objects and
|
||||||
will implicitly refer to objects and arrows in this category.
|
arrows will implicitly refer to objects and arrows in this category.
|
||||||
%
|
%
|
||||||
\subsection{Monoidal formulation}
|
\subsection{Monoidal formulation}
|
||||||
The monoidal formulation of monads consists of the following data:
|
The monoidal formulation of monads consists of the following data:
|
||||||
|
@ -1226,7 +1253,7 @@ they range over are universally quantified.
|
||||||
|
|
||||||
\subsection{Kleisli formulation}
|
\subsection{Kleisli formulation}
|
||||||
%
|
%
|
||||||
The kleisli-formulation consists of the following data:
|
The Kleisli-formulation consists of the following data:
|
||||||
%
|
%
|
||||||
\begin{align}
|
\begin{align}
|
||||||
\label{eq:monad-kleisli-data}
|
\label{eq:monad-kleisli-data}
|
||||||
|
@ -1262,18 +1289,18 @@ This data must satisfy:
|
||||||
%
|
%
|
||||||
Here likewise the arrows $f \tp \Arrow\ X\ (\EndoR\ Y)$ and $g \tp
|
Here likewise the arrows $f \tp \Arrow\ X\ (\EndoR\ Y)$ and $g \tp
|
||||||
\Arrow\ Y\ (\EndoR\ Z)$ are universally quantified (as well as the objects they
|
\Arrow\ Y\ (\EndoR\ Z)$ are universally quantified (as well as the objects they
|
||||||
range over). $\fish$ is the kleisli-arrow which is defined as $f \fish g \defeq
|
range over). $\fish$ is the Kleisli-arrow which is defined as $f \fish g \defeq
|
||||||
f \rrr (\bind\ g)$ . (\TODO{Better way to typeset $\fish$?})
|
f \rrr (\bind\ g)$ . (\TODO{Better way to typeset $\fish$?})
|
||||||
|
|
||||||
\subsection{Equivalence of formulations}
|
\subsection{Equivalence of formulations}
|
||||||
%
|
%
|
||||||
In my implementation I proceede to show how the one formulation gives rise to
|
In my implementation I proceed to show how the one formulation gives rise to
|
||||||
the other and vice-versa. For the present purpose I will briefly sketch some
|
the other and vice-versa. For the present purpose I will briefly sketch some
|
||||||
parts of this construction:
|
parts of this construction:
|
||||||
|
|
||||||
The notation I have chosen here in the report
|
The notation I have chosen here in the report
|
||||||
overloads e.g. $\pure$ to both refer to a natural transformation and an arrow.
|
overloads e.g. $\pure$ to both refer to a natural transformation and an arrow.
|
||||||
This is of course not a coincidence as the arrow in the kleisli formulation
|
This is of course not a coincidence as the arrow in the Kleisli formulation
|
||||||
shall correspond exactly to the map on arrows from the natural transformation
|
shall correspond exactly to the map on arrows from the natural transformation
|
||||||
called $\pure$.
|
called $\pure$.
|
||||||
|
|
||||||
|
@ -1283,7 +1310,7 @@ In the monoidal formulation we can define $\bind$:
|
||||||
\bind \defeq \join \lll \fmap\ f
|
\bind \defeq \join \lll \fmap\ f
|
||||||
\end{align}
|
\end{align}
|
||||||
%
|
%
|
||||||
And likewise in the kleisli formulation we can define $\join$:
|
And likewise in the Kleisli formulation we can define $\join$:
|
||||||
%
|
%
|
||||||
\begin{align}
|
\begin{align}
|
||||||
\join \defeq \bind\ \identity
|
\join \defeq \bind\ \identity
|
||||||
|
|
|
@ -1,7 +1,7 @@
|
||||||
\chapter{Introduction}
|
\chapter{Introduction}
|
||||||
Functional extensionality and univalence is not expressible in
|
Functional extensionality and univalence is not expressible in
|
||||||
\nomen{Intensional Martin Löf Type Theory} (ITT). This poses a severe limitation
|
\nomen{Intensional Martin Löf Type Theory} (ITT). This poses a severe limitation
|
||||||
on both i. what is \emph{provable} and ii. the \emph{reusability} of proofs.
|
on both i. what is \emph{provable} and ii. the \emph{re-usability} of proofs.
|
||||||
Recent developments have, however, resulted in \nomen{Cubical Type Theory} (CTT)
|
Recent developments have, however, resulted in \nomen{Cubical Type Theory} (CTT)
|
||||||
which permits a constructive proof of these two important notions.
|
which permits a constructive proof of these two important notions.
|
||||||
|
|
||||||
|
@ -54,7 +54,7 @@ not true.} There is no way to construct a proof asserting the obvious
|
||||||
equivalence of $f$ and $g$ -- even though we can prove them equal for all
|
equivalence of $f$ and $g$ -- even though we can prove them equal for all
|
||||||
points. This is exactly the notion of equality of functions that we are
|
points. This is exactly the notion of equality of functions that we are
|
||||||
interested in; that they are equal for all inputs. We call this
|
interested in; that they are equal for all inputs. We call this
|
||||||
\nomen{pointwise equality}, where the \emph{points} of a function refers
|
\nomen{point-wise equality}, where the \emph{points} of a function refers
|
||||||
to it's arguments.
|
to it's arguments.
|
||||||
|
|
||||||
In the context of category theory functional extensionality is e.g. needed to
|
In the context of category theory functional extensionality is e.g. needed to
|
||||||
|
@ -73,7 +73,7 @@ The proof obligation that this satisfies the identity law of functors
|
||||||
\end{align*}
|
\end{align*}
|
||||||
%
|
%
|
||||||
One needs functional extensionality to ``go under'' the function arrow and apply
|
One needs functional extensionality to ``go under'' the function arrow and apply
|
||||||
the (left) identity law of the underlying category to proove $\idFun \comp g
|
the (left) identity law of the underlying category to prove $\idFun \comp g
|
||||||
\equiv g$ and thus close the goal.
|
\equiv g$ and thus close the goal.
|
||||||
%
|
%
|
||||||
\subsection{Equality of isomorphic types}
|
\subsection{Equality of isomorphic types}
|
||||||
|
@ -81,7 +81,7 @@ the (left) identity law of the underlying category to proove $\idFun \comp g
|
||||||
Let $\top$ denote the unit type -- a type with a single constructor. In the
|
Let $\top$ denote the unit type -- a type with a single constructor. In the
|
||||||
propositions-as-types interpretation of type theory $\top$ is the proposition
|
propositions-as-types interpretation of type theory $\top$ is the proposition
|
||||||
that is always true. The type $A \x \top$ and $A$ has an element for each $a :
|
that is always true. The type $A \x \top$ and $A$ has an element for each $a :
|
||||||
A$. So in a sense they have the same shape (greek; \nomen{isomorphic}). The
|
A$. So in a sense they have the same shape (Greek; \nomen{isomorphic}). The
|
||||||
second element of the pair does not add any ``interesting information''. It can
|
second element of the pair does not add any ``interesting information''. It can
|
||||||
be useful to identify such types. In fact, it is quite commonplace in
|
be useful to identify such types. In fact, it is quite commonplace in
|
||||||
mathematics. Say we look at a set $\{x \mid \phi\ x \land \psi\ x\}$ and somehow
|
mathematics. Say we look at a set $\{x \mid \phi\ x \land \psi\ x\}$ and somehow
|
||||||
|
@ -114,6 +114,7 @@ formulate this requirement within our formulation of categories by requiring the
|
||||||
\emph{categories} themselves to be univalent as we shall see.
|
\emph{categories} themselves to be univalent as we shall see.
|
||||||
|
|
||||||
\section{Context}
|
\section{Context}
|
||||||
|
\label{sec:context}
|
||||||
%
|
%
|
||||||
The idea of formalizing Category Theory in proof assistants is not new. There
|
The idea of formalizing Category Theory in proof assistants is not new. There
|
||||||
are a multitude of these available online. Just as a first reference see this
|
are a multitude of these available online. Just as a first reference see this
|
||||||
|
@ -131,7 +132,7 @@ implementations of category theory in Agda:
|
||||||
A formalization in Agda with univalence and functional extensionality as
|
A formalization in Agda with univalence and functional extensionality as
|
||||||
postulates.
|
postulates.
|
||||||
\item
|
\item
|
||||||
\url{https://github.com/pcapriotti/agda-categories}
|
\url{https://github.com/HoTT/HoTT/tree/master/theories/Categories}
|
||||||
|
|
||||||
A formalization in Coq in the homotopic setting
|
A formalization in Coq in the homotopic setting
|
||||||
\item
|
\item
|
||||||
|
@ -160,11 +161,11 @@ canonical form.
|
||||||
|
|
||||||
Another approach is to use the \emph{setoid interpretation} of type theory
|
Another approach is to use the \emph{setoid interpretation} of type theory
|
||||||
(\cite{hofmann-1995,huber-2016}). With this approach one works with
|
(\cite{hofmann-1995,huber-2016}). With this approach one works with
|
||||||
\nomen{extensionals sets} $(X, \sim)$, that is a type $X \tp \MCU$ and an
|
\nomen{extensional sets} $(X, \sim)$, that is a type $X \tp \MCU$ and an
|
||||||
equivalence relation $\sim \tp X \to X \to \MCU$ on that type. Under the setoid
|
equivalence relation $\sim \tp X \to X \to \MCU$ on that type. Under the setoid
|
||||||
interpretation the equivalence relation serve as a sort of ``local''
|
interpretation the equivalence relation serve as a sort of ``local''
|
||||||
propositional equality. This approach has other drawbacks; it does not satisfy
|
propositional equality. This approach has other drawbacks; it does not satisfy
|
||||||
all propositional equalites of type theory (\TODO{Citation needed}), is
|
all propositional equalities of type theory (\TODO{Citation needed}), is
|
||||||
cumbersome to work with in practice (\cite[p. 4]{huber-2016}) and makes
|
cumbersome to work with in practice (\cite[p. 4]{huber-2016}) and makes
|
||||||
equational proofs less reusable since equational proofs $a \sim_{X} b$ are
|
equational proofs less reusable since equational proofs $a \sim_{X} b$ are
|
||||||
inherently `local' to the extensional set $(X , \sim)$.
|
inherently `local' to the extensional set $(X , \sim)$.
|
||||||
|
|
|
@ -43,7 +43,6 @@
|
||||||
\researchgroup{Programming Logic Group}
|
\researchgroup{Programming Logic Group}
|
||||||
\bibliographystyle{plain}
|
\bibliographystyle{plain}
|
||||||
|
|
||||||
\addtocontents{toc}{\protect\thispagestyle{empty}}
|
|
||||||
%% \newtheorem{prop}{Proposition}
|
%% \newtheorem{prop}{Proposition}
|
||||||
\makeatletter
|
\makeatletter
|
||||||
\newcommand*{\rom}[1]{\expandafter\@slowroman\romannumeral #1@}
|
\newcommand*{\rom}[1]{\expandafter\@slowroman\romannumeral #1@}
|
||||||
|
@ -52,13 +51,15 @@
|
||||||
|
|
||||||
\pagenumbering{roman}
|
\pagenumbering{roman}
|
||||||
\maketitle
|
\maketitle
|
||||||
|
\addtocontents{toc}{\protect\thispagestyle{empty}}
|
||||||
\tableofcontents
|
\tableofcontents
|
||||||
\pagenumbering{arabic}
|
\pagenumbering{arabic}
|
||||||
%
|
%
|
||||||
\input{introduction.tex}
|
\input{introduction.tex}
|
||||||
\input{cubical.tex}
|
\input{cubical.tex}
|
||||||
\input{implementation.tex}
|
\input{implementation.tex}
|
||||||
|
\input{discussion.tex}
|
||||||
|
\input{conclusion.tex}
|
||||||
|
|
||||||
\nocite{cubical-demo}
|
\nocite{cubical-demo}
|
||||||
\nocite{coquand-2013}
|
\nocite{coquand-2013}
|
||||||
|
|
|
@ -36,6 +36,8 @@
|
||||||
|
|
||||||
\usepackage{lmodern}
|
\usepackage{lmodern}
|
||||||
|
|
||||||
|
\usepackage{enumerate}
|
||||||
|
|
||||||
\usepackage{fontspec}
|
\usepackage{fontspec}
|
||||||
\usepackage[light]{sourcecodepro}
|
\usepackage[light]{sourcecodepro}
|
||||||
%% \setmonofont{Latin Modern Mono}
|
%% \setmonofont{Latin Modern Mono}
|
||||||
|
@ -52,3 +54,5 @@
|
||||||
\usepackage{unicode-math}
|
\usepackage{unicode-math}
|
||||||
|
|
||||||
%% \RequirePackage{kvoptions}
|
%% \RequirePackage{kvoptions}
|
||||||
|
|
||||||
|
\usepackage{pgffor}
|
||||||
|
|
Loading…
Reference in a new issue